Data path and placement optimization in an integrated circuit through use of sequential timing information

ABSTRACT

A method is provided that includes: determining a minimum clock cycle that can be used to propagate a signal about the critical cycle in a circuit design; wherein the critical cycle is a cycle in the design that has a highest proportionality of delay to number of registers; determining for a circuit element in the circuit design, sequential slack associated with the circuit element; wherein the sequential slack represents a minimum delay from among respective maximum delays that can be added to respective structural cycles of which the circuit element is a constituent, based upon the determined limit upon clock cycle duration; using the sequential slack to ascertain sequential optimization based design flexibility throughout multiple stages of a design flow.

CROSS REFERENCE TO RELATED APPLICATIONS

This application is related to the following commonly-owned, pendingpatent applications filed on even date herewith: U.S. patent applicationSer. No. 11/743,279 entitled Optimizing Integrated Circuit DesignThrough Balanced Combinational Slack Plus Sequential Slack; U.S. patentapplication Ser. No. 11/743,301 entitled Optimizing Integrated CircuitDesign Through Use Of Sequential Timing Information; U.S. patentapplication Ser. No. 11/743,326 entitled Reducing Critical Cycle DelayIn An Integrated Circuit Through Use Of Sequential Slack

BACKGROUND OF THE INVENTION

1. Field of the Invention

The invention relates to designing integrated circuits and moreparticularly to optimization of circuit designs through sequentialtiming information.

2. Description of the Related Art

Traditional RTL-to-Layout synthesis has almost exclusively been focusedon non-sequential optimization techniques. A design methodology based onfixed register and latch positions that uses a combination of statictiming analysis, combinational synthesis, and formal equivalencechecking generally provides an effective decomposition of the overalldesign problem into orthogonal issues which can be dealt withindependently. In conjunction with techniques such as a standard-cellplace-and-route approach, zero-skew clock distribution, and full orpartial scan testing, such a methodology has proven to be robust andpredictable, which stimulated broad adoption of the methodology.Complementary to the methodology's appeal from a modeling perspective,this partitioning also supports independent implementation and marketingof the individual components as point tools for which standardizedinterfaces provide the flexibility to adapt the flow for differentneeds.

In recent years, the boundaries between different design stages havebecome more blurred and increasingly integrated approaches have begun todominate parts of the design methodology. This development is driven inpart by the fact that a performance and area compromise made by avertically partitioned flow does not scale proportionally for increasingchip sizes and shrinking device structures. For example, traditionally,logic synthesis and physical design have been performed by separatetools. In the past, a compact model for a logic function and delaycalculation of the individual gates was sufficient to estimate andcontrol the desired design characteristics in both applications.However, the growing chip density and size make abstract modeling ofphysical effects increasingly difficult. The evolution of a moreintegrated approach to logic and physical synthesis addresses, amongother problems, the growing dominance of wire delays for the overallsystem performance.

Sequential optimization (SO) techniques have been researched for manyyears and there are a number of efficient approaches available that areapplicable to practical designs. SO has the potential to significantlyimprove the performance, area, and power consumption of a circuitimplementation to a degree that is not achievable with traditionalcombinational synthesis methods. SO techniques pertain in general to thestaging and timing of provision of signals to combinational logic pathsin a circuit design.

Sequential synthesis methods of practical interest are retiming, C.Leiserson and J. Saxe, “Optimizing synchronous systems,” Journal of VLSIand Computer Systems, vol. 1, pp. 41-67, January 1983; C. Leiserson andJ. Saxe, “Retiming synchronous circuitry,” Algorithmica, vol. 6, pp.5-35, 1991, and clock skew scheduling J. P Fishburn, “Clock skewoptimization,” IEEE Transactions on Computers, vol. 39, pp. 945-951,July 1990. In both methods, the traditionally stated goal is to balancethe path delays between registers and thus to maximize the performanceof the design without changing its input/output behavior.

Retiming is a structural transformation that moves a registers in acircuit without changing the positions of the combinational gates.Changing register positions has the effect of changing the staging ofsignals in a design that is changing the location in the design relativeto combinational logic paths in the design, at which signals aretemporarily stored (i.e. registered). Of course, changing registerpositions also has the effect of changing the timing of the provision ofsignals to combinational logic paths in the design. It is traditionallylimited to designs with extremely high performance requirements, whichare typically developed using a carefully crafted design andverification environment. One reason for a lack of widespread adoptionof retiming has been an inability to accurately predict the impact ofretiming at early design flow stages upon downstream design flow stages.

Clock skew scheduling preserves the circuit structure, but applies tuneddelays to the register clocks—thus virtually moving them in time. Clockskew scheduling—or clock latency scheduling—has the effect of changingthe timing of the provision of signals to combinational logic paths inthe design. In recent years, clock skew scheduling has been adopted insome design flows as a post-layout optimization technique to reduce thecycle time, I. S. Kourtev and E. G. Friedman, Timing Optimizationthrough Clock Skew Scheduling. Boston, Dortrecht, London: KluwerAcademic Publisher, 2000, and the number of close-to-critical paths, C.Albrecht, B. Korte, J. Schietke, and J. Vygen, “Cycle time and slackoptimization for VLSI-chips,” in Digest of Technical Papers of the IEEEInternational Conference on Computer-Aided Design, pp. 232-238, November1999.

SO techniques have found only limited acceptance in contemporary ASICdesign flows, e.g. in pipeline-retiming during logic synthesis or clockskew scheduling for post-layout cycle-time improvement. SO techniquessometimes have been employed as point applications. That is, SOtechniques have been applied to some point in a design flow rather thanto the overall design. For instance, previously published work suggestsretiming as a repeated point-application interleaved with combinationalsynthesis (e.g. S. Malik, E. M. Sentovich, R. K. Brayton, and A.Sangiovanni-Vincentelli, “Retiming and resynthesis: Optimizingsequential networks with combinational techniques,” IEEE Transactions onComputer-Aided Design, vol. 10, pp. 74-84, January 1991). However,computing suitable register positions is computationally expensive andactually not needed in early synthesis stages. Furthermore, such limitedusage may commit a design to a specific structure using only localinformation and missing a significant fraction of the sequentialoptimization space. An analysis of some ASIC designs suggests thatalthough ordinarily only a small fraction of a given design is likely tobe sequentially critical, combinational timing analysis approaches donot reveal this information and thus traditional integrated circuitdesign techniques cannot take advantage of it.

Sequential optimization (SO) techniques have significant potential toimprove the performance of integrated circuit designs and/or decreasetheir size and power consumption. Despite a rich set of theoretical andpractical work on retiming and clock skew scheduling, there has been aneed for improved sequential timing analysis to ascertain an improvedmeasure of sequential flexibility (i.e. sequential slack) in order toachieve better design optimization. There also has been a need forimprovement in the application of sequential flexibility to circuitdesign flow. The present invention meets these needs.

SUMMARY OF THE INVENTION

An aspect of the invention includes obtaining a limit upon clock cycleduration. A determination is made for a circuit element in the circuitdesign, from structural cycles in which the circuit element is aconstituent, as to a limit upon addition of structural cycle delay,based upon a clock cycle duration that is not greater than the obtainedlimit upon clock cycle duration. Therefore, in accordance with an aspectof the invention, a limit upon sequential slack is determined based uponadditional delay that can be added in relation to a circuit elementconsistent with the obtained limit upon cycle duration.

The limit upon clock cycle duration can be made, for instance, byidentifying a structural cycle in a circuit design that has a highestproportionality of structural cycle delay to number of registers. Thisstructural cycle is referred to as the critical cycle. A determinationis made of a minimum clock cycle duration (i.e. maximum clock rate) thatcan be used to propagate a signal about that critical structural cyclewithout incurring a timing violation such as a setup time violation, forexample.

Another aspect of the invention involves determining a first measurementof sequential slack associated with a first portion of a circuit design.The first portion of the circuit design may include structural cyclesthat share a particular circuit element, for example. The firstmeasurement of sequential slack is used to evaluate a first designoptimization choice for the first portion of the circuit design in thecourse of a first stage of an automated circuit design process. Adetermination is made of a second measurement of sequential slackassociated with a second portion of the circuit design. The secondportion of the circuit design may include structural cycles that share adifferent particular circuit element, for example. The secondmeasurement of sequential slack is used to evaluate a second designoptimization choice for the second portion of the design in the courseof a second stage of the automated circuit design process. Sequentialtiming within the circuit design is adjusted. Adjustment of circuittiming occurs after completion of the first stage and after commencementof the second stage. In some embodiments, the first stage is a part of‘logical’ flow and the second stage is a part of the ‘physical’ flow.For instance, the first stage may involve one or more operations suchas, RTL source optimization, data path optimization or logicoptimization, and the second stage may involve one or more operationssuch as, sequential placement, in-place optimization or post-placementretiming. Alternatively, the first stage may occur later in the designflow than the second stage. For instance, routing, which ordinarilyoccurs during the physical flow, could affect either one or both ofplacement or logic optimization in accordance with embodiments of theinvention. Moreover, the first and second stages may involve operationswithin the overall logic stage or within the overall physical stage. Forexample, the first stage might comprise data path optimization and thesecond stage might comprise logic optimization. Adjustment of sequentialtiming may include one or more of clock latency scheduling or retiming,for example. Accordingly, sequential slack can be applied in the courseof multiple stages of a circuit design process.

Yet another aspect of the invention involves determining sequentialslack associated with a respective register in a circuit design. Thedetermined sequential slack is used to evaluate a first design choiceinvolving a structural cycle in the circuit design in which the registeris a constituent. In some embodiments, the first design choice mayinvolve one or more of RTL source optimization, data path optimizationor logic optimization, for example. A maximum critical cycle delaywithin the circuit design is determined. The maximum critical cycledelay is used to select a second design choice in the circuit designthat decreases the maximum critical cycle delay. In some embodiments,the second design choice may involve one or more of placement choices orpost placement retiming choices. Hence, sequential slack is used to addto flexibility of design choices, and a critical cycle can be used as aguide to reduction in maximum cycle time.

An additional aspect of the invention involves obtaining a limit uponclock cycle duration applicable to the circuit design. A determinationis made for each of multiple respective circuit elements in the circuitdesign, of a respective limit upon addition of structural cycle delaythat is not greater than each individual maximal limit of structuralcycle delay for each individual structural cycle in which suchrespective circuit element is a constituent, based upon a clock cycleduration that is not greater than the obtained first limit upon clockcycle duration. A visual representation of the circuit design isprovided. In some embodiments the visual representation is presented ascomputer program code, and in others, it is presented as a computergenerated schematic diagram. The visual representation is annotated toindicate the determined limits upon addition of structural cycle delayfor at least one circuit element shown in the visual representation.Therefore, a circuit designer can take sequential optimization intoaccount in optimizing a design at the RTL stage.

In a still further aspect of the invention, a limit upon clock cycleduration is obtained. A determination is made for a register in thecircuit design, from structural cycles in which the circuit element is aconstituent, of a limit upon addition of structural cycle delay to suchstructural cycles, assuming a clock cycle duration that is not greaterthan the obtained limit upon clock cycle duration. Combinational slackis determined for a combinational path in the circuit design. Thedetermined limit upon addition cycle delay and the determinedcombinational slack are used to identify a change in a data path in thecircuit design involving both a combinational logic change and registerretiming. The identified change is made to the design. By usingtogether, both a measure of sequential flexibility and a measure ofcombinational slack, design changes involving both retiming andcombinational logic changes can be identified and implemented morereadily.

A further aspect of the invention involves obtaining a limit upon clockcycle duration applicable to the circuit design. A determination is madefor each of multiple respective circuit elements in the circuit design,of a respective limit upon addition of structural cycle delay that isnot greater than each individual maximal limit of structural cycle delayfor each individual structural cycle in which such respective circuitelement is a constituent, based upon a clock cycle duration that is notgreater than the obtained limit upon clock cycle duration. Adetermination is made of combinational slack of multiple paths in thecircuit design. Both the determined limits upon addition of cycle delayand the determined combinational slack are used to evaluate multiplepossible combinational logic changes to the circuit design. At least oneof the changes is made. By using together, both a measure of sequentialflexibility and a measure of combinational slack, design changesinvolving logic optimization can be identified and implemented morereadily.

An additional aspect of the invention involves determining at leastfirst and second register placement alternatives. A determination ismade for the first register placement alternative, of a first valueindicative of a proportionality of delay to number of registers for astructural cycle of the first placement alternative having a largestproportionality of delay to number of registers. A determination is madefor the second register placement alternative, of a second valueindicative of a proportionality of delay to number of registers for astructural cycle of the second placement alternative having a largestproportionality of delay to number of registers. The determined firstand second values are used to evaluate the first and second registerplacement alternatives. One of the first and second register placementalternatives is selected. Therefore, a retiming alternative having alesser minimum cycle time can be selected. Moreover, if one of theregister placement alternatives involves a critical cycle within thecircuit design and the other does not, then selection of the registerplacement alternative with the lesser minimum cycle time can have theeffect of reducing the minimum cycle delay of the overall circuit.

A further aspect of the invention involves a method used in defining aclock tree network in a circuit design. First and second non-commonclock tree paths are identified in the design that are used to impart aclock signal to trigger first and second registers that are constituentsof a structural cycle. Path delay from the first register to the secondregister is determined. Path delay from the second register to the firstregister is determined. Based upon a target clock period, adetermination is made of a change in delay in one or both of the firstand second non-common clock tree paths that increases permissiblevariation of clock accuracy, consistent with the target clock period.Delay of one or both of the first and second non-common clock tree pathsis changed according to the determination so as to achieve increasedpermissible variation of clock accuracy. Thus, sequential slack can beused advantageously to relax the accuracy required in creating a clocktree network.

These and other features and advantages of the invention will becomemore apparent from the following description of embodiments thereof inconjunction with the drawings.

BRIEF DESCRIPTION OF THE DRAWINGS

FIG. 1 is an illustrative diagram of a first design flow for the designof integrated circuit in accordance with some embodiments of theinvention.

FIG. 2 is an illustrative diagram of a second design flow for the designof an integrated circuit in accordance with some embodiments of theinvention.

FIGS. 3A-3C are illustrative drawings of a portion of a circuit design(FIG. 3A), the circuit design annotated with timing for zero clock skew(FIG. 3B) and the circuit design annotated with timing for a prescribedclock skew scheduling (FIG. 3C).

FIG. 4 is an illustrative drawing of the circuit design of FIG. 3A withclock period retiming.

FIGS. 5A-5B are illustrative drawing of a portion of the circuit designof FIG. 3A with two modules spaced far apart (FIG. 5A) and with acombination of both clock skew scheduling and clock period retiming(FIG. 5B).

FIGS. 6A-6B are an illustrative drawings of an example circuit (FIG. 6A)and a register timing graph (FIG. 6B) for the circuit. Each node of thegraph of FIG. 6B represents a register of FIG. 6A.

FIG. 7 is an illustrative drawing of a register timing graph for thecircuit of FIG. 6A with clock skew scheduling and with critical pathhighlighted.

FIG. 8 is an illustrative drawing of the register timing graph of FIG. 7with alternative latencies.

FIG. 9 is an illustrative drawing of an alternative register timinggraph for the example circuit of FIG. 6A, that demonstrates that“ordinary sequential slack” in accordance with a prior approach tosequential slack analysis does not reflect sequential flexibility.

FIGS. 10A-10D are illustrative drawings of transformations of a registertiming graph for the circuit of FIG. 6B in accordance with a balancedcombinational slack algorithm used in some embodiments of the invention.

FIGS. 11A-11C are illustrative register timing graphs showing examplesof an original unbalanced graph (FIG. 11A), slack distribution balancedover all edges (FIG. 11B) and slack distributed over three edges with afourth edge handled as a simple constraint (FIG. 11C).

FIGS. 12A-12C are illustrative diagrams showing a register-to-registergraph with delays on the edges (FIG. 12A), a diagram showing truesequential slack for each node (FIG. 12B) and a diagram showing themaximal compatible slack (FIG. 12C).

FIGS. 13A-13D are illustrative diagrams showing a register transferdiagram (FIG. 13A), a diagram showing true sequential slack of nodes xand y (FIG. 13B), a diagram showing maximal compatible slack (FIG. 12C),and a diagram showing balanced compatible slack (FIG. 13D).

FIGS. 14A-14B are illustrative drawings of a partial circuit designbefore (FIG. 14A) and after (FIG. 14B) data path optimization inaccordance with some embodiments of the invention.

FIGS. 15A-15B are illustrative drawings of an example circuit before(FIG. 15A) and after (FIG. 15B) a possible logic optimization that canbe subject to sequential analysis in accordance with some embodiments ofthe invention.

FIG. 16A-16B are illustrative drawings of two candidate placements of acircuit design that can be subject to sequential analysis in accordancewith some embodiments of the invention.

FIG. 17A-17B are illustrative drawings of two candidate placements of acircuit design that can be subject to post-placement sequential analysisin accordance with some embodiments of the invention.

FIG. 18 is an illustrative drawing showing a portion of a circuit thatincludes two example registers coupled in a loop and showing non-commonportions of a clock tree network coupled to the registers and that canbe subject to sequential analysis in accordance with some embodiments ofthe invention.

FIG. 19A-19B shows the distribution of slack over all theregister-to-register paths in an example typical design (paths with over500 ps of slack were omitted) (FIG. 19A) and the distribution of slacksfor the same circuit, after clock latencies have been determined for allregisters using the balanced combinational slack algorithm describedwith reference to FIGS. 10A-10-D (FIG. 19B).

FIG. 20 is an illustrative table that shows the characteristics of thedifferent sequential logic synthesis test cases and the results weobtained.

FIG. 21A-21D are illustrative representations of four differentplacements using different placement techniques. —FIG. 21A shows aplacement obtained using a traditional quadratic programming-basedplacement tool, similar to GORDIAN; FIG. 21B shows the same placementobtained using a combinational slack-driven placer, using similartechniques to those found in modern timing-driven placement tools; FIG.21C shows the design placed using CAPO, a leading-edge placement tooldeveloped at UCLA; and FIG. 21D shows a placement obtained using aprototype sequential placement tool developed in accordance with someaspects of the present invention.

FIGS. 22A-22D show a progression of placement results using theEncounter™ tool and the balanced combinational slack algorithm appliedto one example circuit.

FIG. 23 is a table showing experimental results for a placement tool inaccordance with some aspects of the invention, on a set of tenindustrial benchmark circuits as well as five of the largest synchronousdesigns from the ISCAS89 benchmark suite.

FIG. 24 is an illustrative block level diagram of a computer system thatcan be programmed to implement processes involved with the optimizationof circuit design using sequential timing information in accordance withembodiments of the invention.

DETAILED DESCRIPTION OF THE PREFERRED EMBODIMENTS

The following description is presented to enable any person skilled inthe art to make and use a system and method for sequential timinganalysis and for using sequential optimization to drive one or morestages in an integrated circuit design process. The descriptiondiscloses the system and method in accordance with embodiments of theinvention and in the context of particular applications and theirrequirements. Various modifications to the preferred embodiments will bereadily apparent to those skilled in the art, and the generic principlesdefined herein may be applied to other embodiments and applicationswithout departing from the spirit and scope of the invention. Moreover,in the following description, numerous details are set forth for thepurpose of explanation. However, one of ordinary skill in the art willrealize that the invention might be practiced without the use of thesespecific details. In other instances, well-known structures andprocesses are shown in block diagram form in order not to obscure thedescription of the invention with unnecessary detail. Thus, the presentinvention is not intended to be limited to the embodiments shown, but isto be accorded the widest scope consistent with the principles andfeatures disclosed herein.

Overview of Sequential Optimization Across IC Design Flow Stages

FIG. 1 is an illustrative diagram of a first design flow 100 for thedesign of an integrated circuit in accordance with some embodiments ofthe invention. The first design flow 100 receives as input a registertransfer level (RTL) description 102 of a circuit design. The designflow includes an overall logic synthesis process stage 104 and anoverall physical optimization process stage 106.

In this example design flow, the overall logic optimization andsynthesis process stage 104 includes an interactive source optimizationstage 104-1, an example of which is described below in the section, ‘RTLSource Optimization Example’. The overall logic optimization andsynthesis stage 104 also includes a data path optimization stage 104-2,an example of which is described in the section, ‘Data Path OptimizationExample’ with reference to FIGS. 14A-14B. The overall logic optimizationand synthesis process stages 104 also includes a logic-optimizationstage 104-3, an example of which is described in the section, ‘LegalOptimization Example’ with reference to FIGS. 15A-15B.

The overall physical optimization processes stage 106 includes asequential placement stage 106-1, an example of which is described inthe section, ‘Placement Example’ with reference to FIGS. 16A-16B. Theoverall physical optimization process 106 includes an in-place logicoptimization stage 106-2, an example of which is described in thesection, ‘Logic Optimization Example’ with reference to FIGS. 15A-15B.The physical optimization stage 106 also includes an optionalpost-placement retiming stage 106-3, an example of which is described inthe section, ‘Post-Placement Retiming Example’ with reference to FIGS.17A-17B. The overall physical optimization stage 106 further includes anintegrated clock scheduling and clock tree definition stage 106-4, anexample of which is described in the section, ‘Clock Tree Example’ withreference to FIG. 18.

A sequential timing analysis process 108 provides information used todrive processes within the overall logic synthesis stage 104 and todrive processes within the overall physical optimization stage 106. Inparticular, the sequential timing analysis process 108 determinesmeasurements of sequential slack associated with registers in thecircuit design. Different registers constitute different portions of thedesign; and different registers may have different sequential slacksassociated with them. Thus, the sequential timing analysis process 108provides sequential timing information at multiple stages of the designprocess, that indicates the degree to which sequential optimization(e.g. through adjustment of synchronization latencies) can be used toachieve design objectives such as minimal clock delay, for instance. Theavailability of sequential optimization to address such timing issues,for example, may influence the manner in which design issues areaddressed at different stages of the design flow. As such, sequentialtiming analysis and the possibility of sequential optimization toaddress timing issues adds additional design freedom at various stagesof the design flow.

Specifically, for example, during the overall logic optimization andsynthesis stage 104, the sequential slack measurements are used asmeasures of design flexibility during the interactive sourceoptimization stage 104-1, data path optimization stage 104-2 and logicoptimization stage 104-3. It will be appreciated that these optimizationstages have the effect of changing combinational logic and path delay,which in turn changes the sequential slack. The sequential timinganalysis process 108 repeatedly updates sequential slack measurements inthe course of the overall logic synthesis stage 104.

Similarly, for example, during the overall physical optimization stage106, the sequential slack measurements are used as measures of designflexibility during the sequential placement stage 106-1, in-place logicoptimization stage 106-2, post-placement retiming stage 106-3 and clocktree definition stage 106-4. These physical optimization stages can havethe effect of altering the sequential slack, and therefore, thesequential timing analysis process 108 repeatedly updates sequentialslack measurements in the course of the overall physical optimizationstage 106.

In some embodiments, actual sequential timing (not shown) is performedafter completion of the overall logic synthesis stage 104. Moreover,sequential timing may be further delayed until after completion of theoverall physical optimization stage 106. That is, fixing of clock skewscheduling and retiming (if retiming is to be performed) is delayeduntil late in the overall design process. An advantage of this delayedsequential timing is that sequential optimization flexibility ismaintained throughout the entire design flow. Since register timing andpositioning are not fixed until late in the design flow, sequentialslack can be continually recalculated and used over-and-over as ameasure of design flexibility throughout much of the design flow. Itwill be understood that sequential timing may be insinuated into thephysical design steps that occur later in the overall design process,such as in the post-placement retiming stage 106-3 and in the clock treedefinition stage 106-4. Thus, delaying synchronization latency decisionsuntil later in the design flow enhances design flexibility by permittingsequential optimization throughout multiple stages of the design flow.

In some embodiments, for example, the sequential timing analysis process108 may use time information developed from one or more of the logicsynthesis stages to drive a subsequently performed physical optimizationstage. Moreover, the sequential timing analysis process 108 may use timeinformation developed from one or more of the sub-stages or operationswithin the logic synthesis stage to drive a different sub-stage oroperation within the logic synthesis stage. Conversely, the sequentialtiming analysis process 108 may use time information developed from oneor more of the sub-stages or operations within the physical optimizationstage to drive a different sub-stage or operation within the physicaloptimization stage.

Furthermore, the process of the illustrative flow diagram of FIG. 1 maybe iterative in that the logic synthesis stage and the physicaloptimization stage 106 recurse. In that situation, time informationdeveloped from a stage in the overall design flow may be used to drivean earlier stage in the overall design flow. For instance, the overalllogic synthesis stage 104 may flow to completion followed by one or moresub-stages of the overall physical optimization stage 106. One or moresub-stages of the overall logic synthesis stage 104 then may berepeated. In that situation, the sequential timing analysis process 108may use time information developed from one or more of the physicaloptimization stages to drive a subsequently performed logic synthesisstage. Alternatively, for example, in an iterative flow, the sequentialtiming analysis process 108 may use time information developed from astage later in the logic synthesis flow to drive a subsequentlyperformed stage earlier in the logic synthesis flow. As yet anotheralternative, for example, in an iterative flow, the sequential timinganalysis process 108 may use time information developed from a stagelater in the physical optimization flow to drive a subsequentlyperformed stage earlier in the physical optimization flow.

Subsequent stages of the design flow 100 entail routing and chipfinishing processes 110.

The processes that comprise the design flow 100 may be implemented usingcomputer software programs often referred to as design tools. Thesetools may run on different computer systems. For instance, the toolsinvolved in the overall logic synthesis stage 104 of the flow may run ona first computer system and may produce a partially completed designencoded in computer readable medium. That partially completed designthen may be loaded into a second computer system that runs toolsinvolved in the physical optimization stage 106 of the flow.

Sequential optimization through clock latencies at the register can beapplied through a combination of retiming and classical clock latencyscheduling. In order to realize different clock latencies by classicalclock latency scheduling, a sophisticated clock tree synthesis typicallyis employed. One possibility to implement large clock latencies is theuse of multiple clock domains. See, K. Ravindran, A. Kuehlmann, and E.Sentovich, “Multi-domain clock skew scheduling,” in Digest of TechnicalPapers of the IEEE/ACM International Conference on Computer-AidedDesign, (San Jose, Calif.), pp. 801-808, November 2003, which isexpressly incorporated herein by this reference.

Often, clock tree synthesis, clock skew scheduling, retiming, andphysical placement are considered as one optimization problem. With sucha combined consideration, the available sequential slack can bedistributed optimally to compensate for the additional delay of longinterconnect, for process variations of gates and wires, and forunintentional skews in the clock distribution. The iterative physicaloptimization stage 106 of FIG. 1 outlines such an integrated approach.Placement, clock skew scheduling, and clock tree definition occur withinan iterative overall physical placement stage 106. Furthermore, anin-place logic optimization stage 106-2 driven by sequential timinganalysis process 108 can be used to locally optimize timing-criticalparts. A post-placement retiming stage 106-3 can optionally be appliedto perform a coarse grain implementation of the synchronizationlatencies. Alternatively, multiple clock domains can be utilized torealize large latency differences.

FIG. 2 is an illustrative diagram of a second design flow 200 for thedesign of an integrated circuit in accordance with some embodiments ofthe invention. The second design flow 200 receives as input an RTLdescription 202 of a circuit design. The flow 200 includes a logicoptimization stage 204 and a physical optimization stage 206. The flowof FIG. 2 shows additional details for the flow of FIG. 1, but also doesnot show some of the intermediate steps. For instance, the logicsynthesis stage 104 of FIG. 1 generally corresponds to the logicoptimization stage 204 of FIG. 2, and the physical optimization stage106 of FIG. 1 generally corresponds to the physical optimization stage206 of FIG. 2.

During the logic optimization stage 204, a sequential timing analysisprocess 204-1 computes the sequential slack associated with circuitelements in the design. Sequential slack provides a measure of therelative sequential design flexibility available in different portionsof the design. Design portions with more sequential slack have moredesign flexibility. Design portions with less sequential slack have lesssequential design flexibility. By sequential design flexibility, it ismeant the degree to which sequential optimization can provideflexibility in design choices, whether the choices involve logicsynthesis or physical optimization, for example.

A logic synthesis process 204-2 involves one or more of the followingoperations: decomposition, extraction, factoring, substitution,reconstitution, elimination, structuring and restructuring, technologymapping, gate sizing, buffer insertion, replication and cloning ofgates. The logic synthesis process 204-2 uses sequential slack inevaluating design alternatives. For instance, in deciding amongdifferent design optimization choices, the design optimization process204-2 factors in sequential slack associated with circuit elements inthe design as one measure of degree of freedom in determining whichchoice to make. In decision step 204-43 a determination is made as towhether improvement in terms of delay, area and/or power continues to bemade. If so, then the processes 204-1, 204-2 and 204-3 repeat. If not,then the flow 200 moves to the physical placement stage 206.

During the physical placement optimization stage 206, in step 206-1 atiming driven initial placement is made using sequential slack. Asequential timing analysis process 206-2 detects a critical cycle (e.g.gates) of a design. In decision step 206-3, a determination is made asto whether cycle delay associated with the identified critical cycle isgreater than the clock period. Performance of the design ordinarily islimited by the largest cycle delay within the design. Moreover, if thereis a cycle delay longer than the target clock period, then this delayshould be reduced in order for the design to meet the target clockperiod. If the delay is greater than the clock period, then in step206-4 an incremental placement is made to shrink critical cycle.Typically in this step, gates in a design are moved to positions suchthat the physical length of one or more cycles is reduced so as toreduce the cycle delay of those cycles. During this step the emphasis ison reducing the cycle delay of portions of the design having thegreatest cycle delay. Next, the sequential timing analysis process 206-2again detects a critical cycle (e.g. gates) of a design. The loopincluding processes 206-2 and 206-4 and decision step 206-3 continuesuntil cycle delay is no longer greater than clock period.

It will be appreciated that physical optimization may involve additionalphysical synthesis stages such as, floorplanning and routing amongothers, for example.

In general, the minimum cycle duration (i.e clock period) achievablethrough clock skew scheduling is determined not only based upon a valueof the maximum proportionality of structural cycle delay to number ofregisters for a structural cycle within the design, but also based upona value of the maximum proportionality of combinational delay to numberof registers for all paths from a primary input to a primary output.

Simply stated, a ‘structural cycle’ is a cycle in the design comprisingregisters and combinational gates. In order to simplify the presentationand the computation of the minimum clock cycle duration (i.e. clockperiod) and sequential slack, a graph representing signal paths withinthe design is transformed as follows: all the nodes corresponding toprimary inputs and primary outputs are combined into one single node.With this transformation only cycles need to be considered to ascertainthe minimal clock cycle duration. As used herein, ‘structural cycle’includes both structural cycles in the design and paths from a primaryinput to a primary output transformed to cycles.

As explained more fully below, as used herein, the critical cycle is astructural cycle in the design having a highest proportionality of pathdelay to number of registers. As such, the critical cycle establish aminimum clock cycle from a SO perspective. Therefore, optimizing thedesign to reduce the proportionality of path delay to number ofregisters in the critical cycle has the effect of reducing the minimumclock cycle.

Once the critical cycle delay has been reduced so as to permit thetarget clock period, the second flow 200 proceeds to a coarse timingimplementation by in-place retiming process 208. Basically, the retimingprocess moves the registers in a placed design to best fit the clocklatencies determined for best performance of the design. In general, itis better to use retiming to move registers with large latencies, asthis makes the subsequent clock tree synthesis process easier toperform. Next, a clock latency scheduling process 210 runs. Finally, thesecond flow 200 reaches a clock tree synthesis routing and chipfinishing processes 212. Clock tree synthesis involves choosing specificbuffers and their manner of interconnection in constructing a clocknetwork that implements the latencies for optimal performance of thedesign.

Note that the retiming process 208 and the clock latency schedulingprocess 210 are deferred so as to maintain sequential optimizationflexibility throughout the logic optimization process stage 204 andthroughout the physical placement optimization stage 206.

Clock Skew Scheduling and Retiming

FIGS. 3A-3C are illustrative drawings of a portion of a circuit design(FIG. 3A), the circuit design annotated with timing for zero clock skew(FIG. 3B) and the circuit design annotated with timing for a prescribedclock skew scheduling (FIG. 3C). FIGS. 3A-3C provide a simple example todemonstrate the effect of clock skew (latency) scheduling and retimingand to illustrate the commonality and differences between them. Forsimplicity, we assume that the setup time and hold time of all registersis 0. See generally, I. S. Kourtev and E. G. Friedman, TimingOptimization through Clock Skew Scheduling, Boston, Dortrecht, London:Kluwer Academic Publisher, 2000, for a discussion of set up and holdtime. Further, let the arrival time of all signals at the inputs ofregisters r₁, r₂, r₃, and r₄ and the other inputs of gate g₃ be 0.Assume the required time for the output of register r₅ is equal to thetarget cycle time of 7 units. FIG. 3B shows the timing diagram for theoriginal circuit with zero skew clocking. The minimum clock period ofthe circuit is determined by its critical path which has a maximum delayof D=14 units from the outputs of registers {r₁, r₂, r₃, r₄} throughgates g₁, g₂, g₃, and g₄ to the input of register r₅.

Clock skew scheduling is based on the concept that a fine-tuned advanceor delay of the clocking of individual registers or latches can movethem virtually in the timing diagram forward or backward, respectively.For example, by advancing the clock signal of registers r₁, r₂, r₃, andr₄, by 5 units (s(r₁,r₂,r₃,r₄)=−5)) these registers are virtually movedforward in the timing diagram. The result is that the effect of thecritical path delay on the cycle time is reduced by 5 units, asindicated in FIG. 3C. The resulting minimum clock period is 9 units. Asimilar result could be achieved by delaying the clock signal ofregister r₅ by 5 units. Note that 5 units is the maximum applicableclock skew in this example. If further increased, the minimum delay ofthe paths from r₁ to r₅ would cause a race condition that may result inan illegal state transition. This race condition is commonly referred toas hold time violation.

FIG. 4 is an illustrative drawing of the circuit design of FIG. 3A withclock period retiming. Retiming can be used to further reduce the cycletime below 9 units. Retiming moves registers physically across gates andcan therefore, potentially alleviate hold time violations. FIG. 4 showsthe result of a minimum clock period retiming for the given circuitexample. In order to balance the delay of the critical cycle, the set ofregisters {r₁,r₂,r₃,r₄} has been relocated and merged to r′₁. Thecritical path and minimum clock period in this configuration is 8 units.Thus, by advancing or delaying the clock signal of r′₁ or r₅,respectively, by one unit, the optimum clock period can be furtherreduced to 7 units.

FIGS. 5A-5B are illustrative drawings of a portion of a circuit designwith two modules 502 and 504 that are spaced far apart before (FIG. 5A)and after (FIG. 5B) a combination of both clock skew scheduling andretiming. FIGS. 5A-5B demonstrate the similarity of clock skewscheduling and retiming but also their differences and how bothtechniques could potentially be combined. The two modules 502, 504 inthe example circuit design are placed far apart with a correspondinglarge wire delay of 20 units. Clock skew scheduling alone is not able tosignificantly reduce the clock period. The hold time constraint betweenregisters {r₃,r₄} and r₅ limits the best achievable clock period to 18units. However, as shown in FIG. 5B, by retiming {r₃,r₄} to {r′₃,r′₄}the clock period can be decreased to 12 units. The clock period can befurther reduced to 10 units, if the clocks of {r₁,r₂} and r₆ are skewednegatively and positively, respectively, each by 2 units.

Sequential Timing Analysis

Sequential timing analysis is used to drive a sequential synthesis flowin accordance with some embodiments of the invention. Sequential timinganalysis includes determining a lower bound for a clock cycle (i.e.minimum clock period) achievable by some combination of clock skewscheduling and retiming. A critical cycle is the structural cyclecomprising registers and gates that determines the minimum clock periodthat can propagate a signal within the circuit without incurring atiming violation, specifically a setup time violation. As such, theminimum clock cycle serves as a clock cycle limit. A critical cycle is astructural cycle, which is among structural cycles of a circuit, thecycle with a maximum value for {Delay=total_delay/num_registers}. Inother words, the critical cycle is the structural cycle within thecircuit with the highest proportionality of delay to number ofregisters.

Sequential slack for a circuit element (whether a logic gate or aregister) is the maximum additional delay that circuit element canassume before it becomes part of the critical cycle. Thus, thesequential slack of a circuit element is determined with reference tothe critical cycle. Note the difference between combinational andsequential slack: the combinational slack is derived from a target clockperiod and the actual path delays of the circuit with fixed registerpositions and clock skews. Sequential slack is based on a lower bound ofa circuit's clock period and reflects the sequential criticality ortemporal mobility of registers and combinational gates assuming that theminimum clock period can be realized by retiming and/or clock skewscheduling. Analogous to the application of combinational slack intraditional logic synthesis, the sequential slack can be used to drive atransformation-based synthesis flow for exploring sequentialrestructuring.

Critical Cycle Computation

FIGS. 6A-6B are an illustrative drawings of an example circuit (FIG. 6A)and a register timing graph (FIG. 6B) for the circuit. Each node of thegraph of FIG. 6B represents a register of FIG. 6A. The graph has an edgefrom one node to another node if a logic path starts and ends at thecorresponding registers associated with each edge is the maximum delayof all paths between the corresponding registers.

In the following discussion we assume that the circuit does not have anyprimary inputs or primary outputs. Should the circuit have primaryinputs or primary outputs, these primary inputs or primary outputs canbe combined into one single node so as to tie each path from primaryinput to primary output within a structural cycle.

If the clock latency at each register of the example circuit of FIG. 6Ais set to be zero, the design can run no faster than a clock period of14. In this case the clock period is constrained by the delay fromregister d to register a.

FIG. 7 is an illustrative drawing of a register timing graph for thecircuit of FIG. 6A with clock skew scheduling and with critical pathhighlighted. The latency of register a is increased by 4, i.e. use delaytuning to cause the clock to arrive 4 time units later than normal. Thelatency of register b can be increased by 3, the latency of register ccan be increased by 4, and the arrival of the clock at register d can beleft unchanged.

With these new latencies, the example circuit is able to run at a clockperiod of 10. One can easily verify that, given this clock period, eachedge in the graph still meets its timing requirements. For example, theedge from d to a still has a delay of 14, but since the clock atregister a is delayed by 4 units, any signal along this edge will stillmeet the setup requirement at register a.

It turns out that for this example the given clock latencies areoptimum: there does not exist another clock schedule which will allowthe circuit to run with a faster clock period. This can be seen this byexamining the cycle with edges (a,b), (b,d), (d,a) in FIG. 7. Theseedges are timing critical, in that the signals arriving on these edgesjust exactly meet their required setup times. As such, there does notexist an assignment of latencies that yields a smaller clock period,since such a smaller clock period would result in one of these edgeshaving a timing violation. These edges thus form the critical cycle forthis circuit design. Every cycle in the register timing graph has a meancycle time, which is defined as the total delay around that cycledivided by the number of registers in that cycle. The critical cycle isthe cycle which has the maximum mean cycle time over all cycles in thegraph.

We define the register timing graph G=(V,E) as follows. Take V to be theregisters of the design, together with an additional vertex v_(ext)representing the primary inputs and primary outputs. Add an edge (u, v)to E if there exists a timing path from u to v in the original circuit.Every edge e=(u, v) is labeled with d(e) the maximum delay between u andv in the circuit. During placement, the estimated wire delays areincluded in d(e).

As noted with the example above, finding the critical cycle isequivalent to computing the maximum mean cycle (MMC) for G, which isgiven by

${\max_{\ell \in C}{\sum\limits_{e \in \ell}\;{{d(e)}/{\ell }}}},$where C is the set of all cycles in G. This MMC equation provides abasis for a more mathematical explanation of a process to computeproportionality of delay to number of registers for the structuralcycles of a design so as to identify a critical structural cycle havingthe highest value for such proportionality and so as to ascertain aminimum clock period that can propagate a signal about that criticalcycle. As such, the MMC is equal to the minimum clock period which maybe obtained using clock skew scheduling. We can use Howard's algorithm,J. Cochet-Terrasson, G. Cohen, S. Gaubert, M. McGettrick, and J.-P.Quadrat, “Numerical computation of spectral elements in max-plusalgebra,” in Proceedings of the IFAC Conference on System Structure andControl, July 1998, (represented below by Algorithm 1) to compute theMMC efficiently. The results of, A. Dasdan, S. S. Irani, and R. K.Gupta, “An experimental study of minimum mean cycle algorithms,” Tech.Rep. UCI-ICS 98-32, University of Illinois at Urbana-Champaign, 1998,suggest that Howard's algorithm is the fastest known algorithm forcomputing MMC.

Algorithm 1 Howard's Algorithm For Computing MMC 1:  for all ^(u ∈ V) do{initial guess} 2:   ^(π(u) ← e for some e = (u,v) ∈ E) 3:  repeat {mainloop} 4:   find all cycles ^(C) ^(π) ^(in G) ^(π)^(= (V,{π(u) : u ∈ V} ⊂ E)) 5:   for all ^(u ∈V) do {in reversetopological order} 6:    if u = CYCLEHEAD^((L) ^(u) ⁾ for some ^(L) ^(u)^(∈ C) ^(π) then 7:     ^(η(u) ←) MEANCYCLETIME^((L) ^(u) ^(),x(u) ← 0)8:    else {defined recursively} 9:     ^(η(u) ← η(v)) where^(π(u) = (u,v)) 10:     ^(x(u) ← x(v) + d(π(u)) − η(v)) 11:   for all^(e = (u,v) ∈ E) do {modify ^(π) } 12:    if ^(η(u) < η(v)) then 13:    ^(π(u) ← e) 14:   if ^(π) unchanged then 15:    for all^(e = (u,v) ∈ E) do 16:     if ^(η(v) = η(u)) and^(x(u) < x(v) + d(e) − η(v)) then 17:      ^(π(u) ← e) 18: until nochange in ^(π) 19: return ^(max) ^(u∈V) ^(η(u))

Generally speaking, the idea behind Howard's algorithm is to maintain asmall set of edges π, which starts as an initial guess of the criticalcycle in the graph. π is chosen such that every vertex has only oneoutgoing edge in the induced subgraph G^(π), which is the graph obtainedby removing all edges except those in π. Now G^(π) contains at least onecycle, and every vertex reaches exactly one cycle through a path inG^(π.)

Every cycle L in G^(π) is associated with an arbitrary distinguishedvertex in L, denoted by CYCLEHEAD(L). Two vertex labels η and x are alsodefined. If L_(v) is the unique cycle reached from v in G^(π), then η(v)is the mean cycle time of L_(v) and x(v) is the sequential delay from vto CYCLEHEAD(L_(v)). Due to the scarcity of π, all cycles of G^(π) canbe enumerated in time linear in the size of the graph; likewise, η(v)and x(v) can be computed in linear time as well.

Given π and the corresponding labels (η,x), Howard's algorithm proceedsby finding a new π which yields new vertex labels which are strictlygreater than the previous labels in a lexicographic sense. This can bedone in a greedy and efficient manner. This procedure is iterated untilno improvement can be made, at which point the critical cycle (and thecorresponding MMC) is identified in π. Some details are omitted here;more information about Howard's algorithm and a proof of correctness canbe found in J. Cochet-Terrasson, et al., “Numerical computation ofspectral elements in max-plus algebra,” Supra.

Sequential Slack

Once the critical cycle has been identified, the sequential criticalityof the vertices of other cycles in the circuit can be determinedrelative to a minimal clock cycle derived from the proportionality ofdelay to number of registers of the critical cycle.

In this section, a prior notion of sequential slack, is presented, whichwas described in, J. Cong and S. K. Lim, “Physical planning withretiming,” in Digest of Technical Papers of the IEEE/ACM InternationalConference on Computer-Aided Design, (San Jose, Calif.), pp. 1-7,November 2000. A significant shortcoming in that prior notion ofsequential slack is demonstrated, and new alternative notion ofsequential slack in accordance with the principles of some embodimentsof this invention is introduced. We refer to that prior notion ofsequential slack as reference vertex dependent sequential slack (RVDsequential slack)

FIG. 8 is an illustrative drawing of the register timing graph of FIG. 7with alternative latencies. Sequential slack gives rise to flexibilityin the implementation of a circuit design. The alternative assignment oflatencies to the registers also meets the circuit timing requirements.More particularly, for example, that the latency of register c haschanged from +4 to +2, independent of the other register latencies,while still meeting the timing requirements. This is what is meant bysaying that register c has sequential flexibility: the timingspecification of register c is not fixed, relative to the critical cycleof the overall design. In contrast, for example, register a issequentially critical, as it lies on the critical cycle and thus itsclock latency cannot be changed without affecting the latencies of theother registers on the critical cycle.

Sequential slack is a measure of degree of flexibility available tooptimize other aspects of a design. As an analogy, for examplecombinational timing analysis provides information that may allowresizing of gates which do not lie on the critical path in order tominimize size, or to introduce longer wires and additional wire delayalong non-critical paths in order to obtain better placements. Likewise,sequential timing analysis can provide information that may allowsimilar trade-offs in the domain of sequential optimization for gatesthat do not lie on a critical cycle.

The work on the prior notion of sequential slack in, J. Cong et al.“Physical planning with retiming,” supra, quantifies the concept ofsequential flexibility through the use of sequential arrival time,sequential required time, and sequential slack. According to that priorwork, given a target clock period of φ the sequential arrival andrequired times at all vertices v ε V with respect to a reference vertexv_(ref) can be computed from Equations (1.1) through (1.3) below, usinga modified version of the Bellman-Ford algorithm, T. H. Cormen, C. E.Leiserson, and R. L. Rivest, Introduction to Algorithms, The MIT Electr.Eng. and Computer Science Series, MIT Press/McGraw Hill, 1990.

$\begin{matrix}{{A_{seq}\left( {v,v_{ref}} \right)} = {{\max\limits_{{({u,v})} \in E}{A_{seq}\left( {u,v_{ref}} \right)}} + {d\left( \left( {u,v} \right) \right)} - \phi}} & (1.1) \\{{R_{seq}\left( {v,v_{ref}} \right)} = {{\min\limits_{{({v,w})} \in E}{R_{seq}\left( {w,v_{ref}} \right)}} - {d\left( \left( {v,w} \right) \right)} + \phi}} & (1.2) \\{{A_{seq}\left( {v_{ref},v_{ref}} \right)} = {{R_{seq}\left( {v_{ref},v_{ref}} \right)} = 0}} & (1.3)\end{matrix}$The sequential slack is S_(seq)=R_(seq)−A_(seq).

A_(seq) and R_(seq) represent respectively the earliest and latestrelative position in time to which a register can be moved (by retimingor clock skewing) while still meeting timing with respect to thereference point. S_(seq) measures the feasible range of temporalpositions for v relative to the reference node. The concepts ofsequential arrival time, required time, and slack are equivalent tothose used in the combinational timing domain.

While this prior definition of sequential slack represents a metric forquantifying sequential criticality, one shortcoming of this priorapproach is that this criticality is only with respect to the givenreference vertex, v_(ref). A different choice of reference vertex,v_(ref) will impose different constraints. Hence, we refer to this priornotion of sequential slack as reference vertex dependent sequentialslack (RVD sequential slack).

FIG. 9 is an illustrative drawing of an alternative register timinggraph for the example circuit of FIG. 6A, that demonstrates that RVDsequential slack in accordance with a prior approach to sequential slackanalysis does not fully reflect sequential flexibility. In FIG. 9, thechoice of v_(ref) gives a non-zero slack value for vertices on thecritical cycle. In this example, vertex c (highlighted) is chosen to bethe reference node v_(ref) when computing the sequential slacks S_(seq).Note that in this example, vertices a, b and d all have positive slack,despite being on the critical cycle. In accordance with principles ofsome embodiments of the invention, there is benefit to having a criticalslack metric which yields a slack of zero for sequentially criticalvertices. This, desire together with the fact that the value forsequential slack according to the above prior work depends on the choiceof reference node, motivates us to conclude that the prior RVD notion ofsequential slack is not optimal for determining sequential criticalityand sequential flexibility.

To overcome the limitations of the prior notion of RVD sequential slack,we have developed a concept which we call reference cycle dependent(RCD) sequential slack (also referred to herein as “true” sequentialslack). RCD sequential slack for a vertex is defined as the maximumdelay which can be added to the outgoing edges of that vertex withoutincreasing the MMC of the design. More specifically, the RCD sequentialslack for a vertex of a structural cycle is the maximum delay that canbe added to outgoing edges of a vertex of that cycle without resultingin that cycle becoming the critical cycle. As such, the MMC serves asthe “reference cycle”. From this definition, it is clear that verticeson the critical cycle will have a RCD (“true”) sequential slack of zero,making this notion of sequential slack a more appropriate measure ofsequential flexibility, compared to ordinary sequential slack.

Computing the true sequential slack of a vertex is straightforward butcan be expensive. One can examine all structural cycles which passthrough the specified vertex, finding the maximum delay which can beadded without violating the MMC, and then taking the minimum of allthese quantities.

More specifically, for example, the RCD sequential slack for a givenregister is determined based upon all structural cycles in which thegiven register is a constituent. The RCD sequential slack for the givenregister is the maximum delay that can be added to a structural cycle ofwhich the given register is a constituent, using a clock cycle at orabout the limit determined based upon the critical cycle, withoutproducing a timing violation, specifically a setup time violation. Inother words, a clock cycle selected for use to determine a maximum delaythat can be added for a given register should be at or about the limitdetermined for the critical cycle. Thus, the determined maximum delay isdetermined based upon such selected clock cycle, which of course, couldbe a clock cycle selected to be at the limit determined based upon thecritical cycle.

Even more specifically, for example, the RCD sequential slack for suchgiven register is the minimum delay from among all respective maximumdelays that can be added to respective structural cycles of which thegiven register is a constituent, based upon a clock cycle at or aboutthe limit determined based upon the critical cycle, without producingsetup time violations.

Thus, the maximum delay that can be added to a structural cycle of whichthe given register is a constituent is the minimum delay from among allrespective maximum delays that can be added to respective structuralcycles of which the given register is a constituent, assuming the use ofa clock cycle at or about the limit determined based upon the criticalcycle, without producing setup time violations.

For instance, assume that the given register is a constituent of threestructural cycles. Also, assume a clock cycle at or about the limitdetermined based upon the critical cycle. Further, assume that a maximumdelay that can be added to a first structural cycle of which the givenregister is a constituent is 3 units; and a maximum delay that can beadded to a second structural cycle of which it is a constituent is 9units; and a maximum delay that can be added to a second structuralcycle of which it is a constituent is 5 units. That is, assuming a clockcycle at or about the determined limit, 3 units of delay can be added tothe first structural cycle without resulting in a setup time violation;up to 9 units of delay can be added to the second structural cyclewithout resulting in a setup time violation; and 5 units of delay can beadded to the third structural cycle without resulting in a setup timeviolation. Then, the RCD sequential slack of the given register is 3units.

A more efficient approach to computing the above RCD (“true”) sequentialslack of a vertex S_(true) is:S _(true)(v)=−max_((u,v)εE) {A _(seq)(u,v)+d(u,v)−φ}The values of A_(seq)(u,v) can be computed with the Belmann-Fordalgorithm. A_(seq)(u,v) is the longest path from v to u with respect tothe length function d(u,v)−φ.

Alternatively, the true sequential slack of any node can be computedusing a combination of the Belman-Ford algorithm and the Dijkstraalgorithm:

First compute a clock schedule with the Belman-Ford algorithm. Let l(u),uεV be the latencies of the nodes. Then the expressionl(u)+d(u,v)−\φ−l(v)is non-positive.

The sequential arrival time A_(seq) (v, v_(ref)) slack for any singleregister can then be computed using Dijkstra's algorithm with the lengthfunctionl(u)+d(u,v)−\φ−l(v)instead of the length functiond(u,v)−\φDijkstra's algorithm can compute longest paths more efficiently if thelengths of the edges are non-positive. The new length function does notchange the length of a cycle and hence the sequential slack valuecomputed. See, R. E. Bellman: On a routing problem. Quarterly of AppliedMathematics 16 (1958), 87-90; E. W. Dijkstra: A note on two problems inconnexion with graphs, Numerische Mathematik 1 (1959), pages 269-271.

As yet another alternative, the Floyd-Warshall algorithm can be used tocompute all-pair shortest paths (length function \φ−d(u,v). See, R. W.Floyd: Algorithm 97—shortest path, Communications of the ACM 5 (1962),345.

Despite this more efficient formulation, computing S_(true) is still acomputationally expensive procedure. An understanding of the concept ofRCD (“true”) sequential slack (as defined herein) is important in atheoretical sense in that the sequential slack metrics used in practicalapplications are intended to be an approximation to the theoreticallyideal true sequential slack.

Balanced Combinational Slack

The complexity of computing RCD sequential slack can make its generaluse throughout a sequential synthesis flow computationally quiteexpensive, where such values might need to be recomputed many times asstructural or physical changes are made to the design. Accordingly, anovel technique has been developed to distribute combinational slack ina generally balanced fashion across a design or portions of a designwhile avoiding the explicit computation of RCD sequential slack. As usedherein combinational slack is the amount by which the delay of a gatecan be increased before a combinational path through the gate becomescritical. One technique to balance path delays is described in commonlyowned co-pending U.S. patent application Ser. No. 11/373,670, filed Mar.10, 2006, entitled OPTIMIZATION OF COMBINATIONAL LOGIC SYNTHESIS THROUGHCLOCK LATENCY SCHEDULING, invented by C. Albrecht, A. Kuehlmann, D.Seibert and S. Richter, which is expressly incorporated herein in itsentirety by this reference.

For each vertex, combinational slack on the edges incident to thatvertex is balanced such that the worst combinational slack among theincoming edges is substantially equal to the worst slack among theoutgoing edges. One technique to achieve such balancing of combinationalslack is to adjust clock latencies applied to registers, represented asvertices in a register transfer diagram, so as to achieve such balancingof combinational slack. This process provides a useful technique fordistributing combinational slack throughout the design, as it helpsensure that optimization tools (e.g. the synthesis techniques describedbelow in the section headed “Sequential Logic Synthesis” are givenflexibility wherever it exists. This process is called the balancedcombinational slack algorithm, which is shown as Algorithm 2.

Algorithm 2 Balanced Combinational Sequential Slack 1:  while graphcontains directed cycle do 2:   find critical cycle C with maximum meancycle time 3:   assign latencies to C which satisfy timing 4:   contractcycle C to a single vertex v_(c) 5:   adjust delays on edges directedinto and out of v_(c) 6:  uncontract graph and compute latencies for alloriginal vertices

The algorithm makes use of the fact that the vertices in a criticalcycle always have their latencies in “lock-step” with each other. Thatis, such vertices have latencies which are all interdependent, andsetting the latency for one vertex fixes the latencies of all the othersin the critical cycle. This is why in each iteration the critical cycleC is contracted into a single vertex v_(c), since only a single latencyneed be computed for the representative (contracted) vertex tocompletely determine the latencies for the entire critical cycle. Thedelays for edges directed into and out of v_(c) are adjusted to accountfor the fact that the latencies assigned to the individual vertices inthe critical cycle may be different from each other. That is, theincoming edges to the new vertex v_(c) may, in fact, have pointed todifferent vertices of C, and this is reflected in the resulting graphafter C is collapsed.

FIGS. 10A-10D are illustrative drawings of transformations of a registertiming graph for the circuit of FIG. 6B in accordance with a balancedcombinational slack algorithm used in some embodiments of the invention.FIG. 10A shows the critical cycle highlighted and having vertices (a, b,d) with latencies of (+1, 0, −3) respectively.

FIG. 10B shows the vertices (a, b, d) of the critical cycle contractedto a single vertex abd. The delays on the edges leading into and out ofabd are adjusted as follows: consider the edge (c, a). Let τ_(v) be thelatency assigned to vertex v. Since a was assigned a latency of +1before being contracted into abd, we can consider this vertex to have alatency of +1 relative to abd, soτ_(a)=1+τ_(abd)   (2.1)Now suppose the latency of c is set so that the edge (c, a) exactlymeets its timing requirements. Thenτ_(c)+10−φ=τ_(a)   (2.2)where φ is the clock period. But in that case the edge (c, abd) mustexactly meet its timing requirement as well, since abd represents a inthe contracted graph. Then we must haveτ_(c) +q−φ=τ _(abd)   (2.3)where q is the delay assigned to the edge (c, abd). Combining Equations2.1 through 2.3 we find q=9.

The procedure of adjusting the edge delays can be easily generalized, sothat edges directed into the contracted vertex have a value subtractedfrom their delays, where this value is the latency of the originalvertex to which the edge was previously directed. Likewise, edgesdirected out of the contracted vertex have the corresponding latencyadded to their delays as shown in FIG. 10B. More particularly, when thealgorithm finds and contracts a cycle the relative differences of thelatencies of the cycle are determined, and the edge delays are adjustedto account for the latencies of vertices subsumed within the contractedgraph. The determination of adjustments is made based upon the followingrelationship representing the slack of an edge (a, b):slack(a, b)=φ−delay(a, b)−τ_(a)+τ_(b)

Referring to FIG. 10A, vertex a has a latency which is by 1 larger thanthe latency of vertex b, and vertex d has a latency which has a latencythat is by 3 smaller than the latency of vertex b. In the modified graphof FIG. 10B-10C, there is only one vertex, the vertex abd, for the threevertices a, b, and d. In the example of FIGS. 10B-10C, the latency ofthe new vertex is equal to the latency of the vertex b in the originalgraph. Thus, the edges in the modified graph of FIGS. 10B-10C (with thecontracted cycle) have a delay such that for each edge the slacks in theoriginal and in the modified graph are equal. Specifically, since theclock on vertex d is advanced by three time units (−3), the delay (6) ondirected edge (abd, c) is reduced by three time units (6−3). Conversely,the delay on (5) on directed edge (c, abd) is increased by three timeunits (5+3). Likewise, since the clock on vertex a is delayed by onetime unit (+1), the delay (10) on directed edge (c, abd) is decreased byone time unit (10−1).

FIG. 10D is an illustrative drawing representing a final step in theprocess in which the graph of FIG. 10C is uncontracted and latencies arecomputed for the original vertices. Since the latency of abd wasassigned a value of +3 and a was given a latency of +1 relative to abd(in the first iteration), then a has an overall latency of +4. Similarcomputations are done for the other vertices. The new latencies provideadditional slack on the edges, which provides additional freedom foroptimization.

The sequential slack indicates how much delay can be added at some pointin a design by changing the latencies at the registers such that setupconstraints are met. Whereas, combinational slack indicates how muchdelay can be added at some point in a design but without changing thelatencies at the registers.

As noted earlier, the final solution obtained has the property that foreach vertex the worst slack of all incoming edges substantially equalsthe worst slack of all outgoing edges, hence we call the algorithmbalanced. In the example, register c has a RCD (“true”) sequential slackof +8, and can take on any latency between +4 and −4. The balancedcombinational slack algorithm assigns register c a latency of 0,balancing the slack so that half of this slack is distributed to theworst outgoing edge (c, a) and the other half to the worst incoming edge(d, c).

More specifically, For a clock period of φ=10, the RCD sequential slackof register c is 8. The combinational slack of the edge (c, a) is 4 andthe combinational slack of the edge (d, c) is also 4 with respect to thelatencies and according to the slack formula, slack(a, b)=φ−delay(a,b)−τ_(a)+τ_(b). If there were additional combinational slack on theedges (a, b) and (b, d), the true sequential slack would be higher.

We note that an equivalent property to the combinational slack balancingproperty is the following global optimization criterion: the vectorobtained by taking the slacks for each edge and sorting them inincreasing order is lexicographically the greatest among all suchvectors. In our example this slack vector is (0, 0, 0, 2, 4, 4, 5); noother assignment of latencies to the registers yields alexicographically greater slack vector. Basically, we computed the slackwith the formula slack (a, b)=φ−delay(a, b)−τ_(a)+τ_(b) for each edge.Edges (a, b), (b, d) and (d, a) have a slack of 0, edge (d, b) has aslack of 2, edges (d, c) and (c, a) have a slack of 4 and edge (c, d)has a slack of 5. Lexicographically greater means that we compare thevector (after sorting) component by component until a component isdifferent, and this first component which is different determines whichof the vectors is lexicographically greater. For example (0, 0, 0, 2, 4,4, 5)>(0, 0, 0, 2, 3, 4, 5), and (0, 0, 0, 2, 4, 4, 5)<(0, 0, 2, 2, 4,4, 5).

Computing RCD sequential slack can be computationally more expensive tocompute that to compute RVD sequential slack or balanced combinationalslack. However, in practice, it is not necessary to compute the RCDsequential slack for all registers (all nodes in theregister-to-register graph). The following approach is possible:

-   1. Compute the balanced combinational slack for all registers of the    design.-   2. Fix the latencies for those registers which have a balanced    combinational slack larger than a specified value.-   3. Compute the RCD sequential slack for the registers which have a    balanced combinational slack less than or equal to the specified    value and for which the latencies have been fixed in step 2.

In the second step, the registers for which the latencies are fixed canbe considered as primary inputs and primary outputs. As mentioned above,primary inputs and primary outputs are combined into a single node forthe computation of sequential slack. Combining the primary inputs andprimary outputs into a single node reduces the size of the graph andhence the complexity of the computation of the sequential slack. Thefinal slack value is either the balanced combinational slack for theregisters for which the latency has been fixed in step 2 or the RCD(“true”) sequential slack computed in step 3. This slack value is alower bound on the RCD sequential slack. If the slack value is smallerthan the specified value used to determine the registers for which thelatency is fixed, then the value is equal to the sequential slack.

The specified slack can be an external parameter set by a user, forexample, or an internal parameter determined so as to achieve anacceptable tradeoff between degree of optimization and runtime of thecomputer software design tool.

The maximum mean cycle (MMC) is the minimum clock period achievable byclock latency scheduling. If the target clock period is equal to thisvalue, all the slacks, the balanced combinational slack and the RCDsequential slack are nonnegative. The balanced combinational slack issmaller than or equal to the RCD sequential slack, and fixing thelatencies for some registers cannot increase the RCD sequential slack.Hence, the slack computed by this scenario is a lower bound of the RCDsequential slack.

Partially Balanced Combinational Slack

In the above subsection it is shown that combinational slack can bedistributed and balanced over all edges. It is also of interest todistribute and maximize the combinational slack only over some of theedges, while the slack of other edges is simply constrained to benonnegative, i.e. the inequality constraint holds.

FIGS. 11A-11C are illustrative register timing graphs showing examplesof an original unbalanced graph (FIG. 11A), slack distribution balancedover all edges (FIG. 11B) and slack distributed over three edges with afourth edge handled as a simple constraint (FIG. 11C). FIG. 11A is anillustrative drawing of a register timing graph corresponding to acircuit (not shown) with unbalanced combinational slack and a clockperiod φ=11. FIG. 11B is an illustrative drawing of a register timingdiagram corresponding to the same circuit as FIG. 11A except withcombinational slack distribution balanced over all four edges. FIG. 11Cis an illustrative drawing of a register timing diagram corresponding tothe same circuit as FIG. 11A except with combinational slack partiallybalanced such that edge (d, c) is simply handled as constraint, i.e.,the edge is required to have nonnegative slack.

An abstract formulation of this problem is as follows: Given a directedgraph G=(V,E), weights c:E→

, and a partition E=E₁∪E₂. Find an assignment τ:V→

such that τ(u)+c((u,v))≦τ(v) for all edges (u,v) ε E₁ and τ(u)+c((u,v))+s ((u,v))≦τ(v) for all edges (u,v) ε E₂, and such that the slackvector

(s((u,v)))_((u,v)-εE) ₂

sorted in non-increasing order is the greatest possible by lexicographicorder. In the following we will refer to the edges in E₁ asnon-parameterized and to the edges in E₂ as parameterized. In theprevious subsection we have solved the problem for E₂=E.

Note, that in this formulation c((u,v)) does not necessarily correspondto the delay of the path from u to v, in fact if the edges correspond toa setup constraint, c((u, v)) would be the delay of the edge minus thegiven clock period. In our example, the variables τ(v), v ε V,correspond to the latencies of the registers. However, in a more generalapplication the nodes can also model other components as we will show inthe next section.

For solving the partial balancing problem, one can modify Howard'salgorithm (Algorithm 1) and solve it in combination with the balancedsequential algorithm (Algorithm 2). For more details the refer to C.Albrecht, B. Korte, J. Schietke, and J. Vygen, “Maximum mean weightcycle in a digraph and minimizing cycle time of a logic chip,” inDiscrete Applied Mathematics, vol. 123, pp. 103-127, November 2002; andC. Albrecht, B. Korte, J. Schietke, and J. Vygen, “Cycle time and slackoptimization for VLSI-chips,” in Digest of Technical Papers of the IEEEInternational Conference on Computer-Aided Design, pp. 232-238, November1999, in which the same basic problem is solved, using a differentalgorithm for the MMC computation.

Examples of RCD (“True”) Sequential Slack, Maximal Compatible SequentialSlack and Balanced Compatible Sequential Slack

FIGS. 12A-12C are illustrative diagrams showing a register-to-registergraph with delays on the edges (FIG. 12A), a diagram showing RCD(“true”) sequential slack for each node (FIG. 12B) and a diagram showingthe maximal compatible slack (FIG. 12C). The register-to-register graphof FIG. 12A is identical to the diagram of FIG. 6B, which represents theexample circuit of FIG. 6A. The critical cycle in the register transferdiagram of FIG. 12A is the structural cycle,a→b→d→a,since (9+7+14)/3=10, which is the maximum average delay per register ofall cycles in the diagram of FIG. 12A. Thus, in the example illustratedin FIG. 12A, the formulation (9+7+14)/3=10 represents a proportionalityof delay to number of registers. Moreover, that formulation representsthe maximum proportionality of delay to number of registers in theexample. Therefore, the minimum clock period, φ=10.

In contrast, note for example, that the structural cycle,c→a→b→d→c,is not critical since it has an average register delay of(10+9+7+6)/4=8, which not the maximum average delay of all cycles underconsideration in the diagram.

In FIG. 12B, the value 8 associated with node c represents the RCDsequential slack. The true sequential slack represents how much delayneeds to be added before or after a given register (in this example aregister represented by node c) such that the register becomes part of acritical cycle. The RCD sequential slack is a value equal to the lengthof the shortest cycle through the node with respect to the lengthfunction φ−d(u,v).

Stated differently, the RCD sequential slack is the minimum delay fromamong respective maximal delays that can be added to respectivestructural cycles in which the a circuit element (e.g. a register)represented by node c is a constituent, based upon use of a clock cycleduration that is not greater than the obtained limit upon clock cycleduration. In this example, node c is a constituent of structural cycle,c→a→b→d→c,and node c is a constituent of structural cycle,c→d→cThe RCD sequential slack for node c calculated as the slack aboutstructural cycle,c→a→b→d→c, is8=(10−10)+(10−9)+(10−7)+(10−6),which means that 8 time units is the greatest delay that can be added tothe corresponding structural cycle without resulting in a setup timeviolation, assuming a clock cycle φ=10.

In contrast, note for example, that RCD sequential slack for thestructural cycle,c→d→c,has a length with respect to φ−d(u,v), which is calculated as,9=(10−5)+(10−6),which means that 9 time units is the greatest delay that can be added tothe corresponding structural cycle without resulting in a setup timeviolation, assuming a clock cycle φ=10.

Since 9 time units is greater than 8 time units, the structural cycle,c→d→c,does not have a length of the shortest cycle through the node withrespect to the length function φ−d(u,v). In other words, the structuralcycle,c→a→b→d→c,has the minimum delay (8) from among respective maximal delays (8) and(9) that can be added to respective structural cycles in which the acircuit element represented by node c is a constituent, assuming a clockcycle φ=10, that is not greater than the obtained limit upon clock cycleduration.

Thus, in this example, adding a delay of 8 before or after the node (orregister) c would make the maximum average delay of all cycles throughnode c be φ=10. For instance, adding a delay of 4 to edge c→a and addinga delay of 4 to edge d→c results in the average register delay for thecycle,c→a→b→d→c being,(10+4)+9+7+(6+4)/4=10.

Thus, in accordance with some embodiments of the invention, based uponuse of a clock cycle φ=10, an RCD sequential slack of 8 time unitsindicates that an additional combinational delay of 8 time units can beadded to the structural cycle,c→a→b→d→c.

This additional combinational delay can be added using the clock cycleφ=10, since timing of a register represented by node c can be adjustedthrough sequential optimization (SO) involving adjustment of clocklatency scheduling or through retiming adjustments, for example.

FIG. 12C is an illustrative drawing showing compatible slack. A set ofcompatible sequential slacks for multiple circuit elements compriseslacks that for any one of the multiple circuit elements, can be used inany way, that for which the other slacks are still valid. For example,assume that a set of registers R₁-R_(N) have a set of compatiblesequential slacks. That is, R₁ has slack S₁, R₂ has slack S₂, . . .R_(N) has slack S_(N). Further, assume, for example, that R1 has asequential slack of 5. In that case, the delay associated with registerR1 can be increased by 1, 2, 3, 4, or 5, and the slacks of all ofregisters R₂-R_(N) are still valid and can be used to their maximum aswell.

The maximal compatible slack represents how much delay should be addedto every register (simultaneously but can be different values fordifferent registers) such that all registers are part of a criticalcycle. Referring to FIG. 12C, for example, nodes (registers) a, b, and deach is part of a critical cycle already, which means that the RCDsequential slack as well as the maximal compatible slack for those nodesis zero. Therefore, additional slack can be added only with respect tonode c, and 8 is the maximal delay value that can be added before orafter node c to make sure that node c is also part of a critical cycle.

FIGS. 13A-13D are illustrative diagrams showing a register transferdiagram (FIG. 13A), a diagram showing RCD sequential slack of nodes xand y (FIG. 13B), a diagram showing maximal compatible slack (FIG. 12C),and a diagram showing balanced compatible slack (FIG. 13D). The cycle,p→u→w→pis critical cycle in the diagram of FIG. 13A. The average delay of thecritical cycle, 10 in this example, is by definition the maximum averagecycle delay for the diagram. The minimum clock period, as determinedfrom the maximum average cycle delay is φ=10.

FIG. 13B shows the RCD sequential slack of nodes x and y to be 10 and 9respectively. The node x is present in one cycle, and the truesequential slack of that one cycle is computed as,10=(10−10)+(10−9)+(10−7)+(10−6)+(10−8).

The node y is present in two cycles. One cycle is the same one that nodex is in. The other is the cycle, y→w→y, which has a RCD sequential slackcomputed as,9=(10−6)+(10−5).

Note that the RCD sequential slacks for nodes x and y cannot be appliedsimultaneously to those nodes. Thus, a determination of maximalcompatible slacks is beneficial. FIG. 13C is an illustrative drawingshowing an example of maximal compatible slack for each of the nodes ofthe diagram of FIG. 13A.

FIG. 13D is an illustrative drawing showing an example of balancedcompatible slack which also is another maximal compatible slack. Abalanced compatible slack is a maximal compatible slack with theadditional property that it is not possible to increase the compatibleslack of one register by decreasing the slack of another register whichhas a larger slack. In contrast, the slacks of FIG. 13C do not representbalanced compatible slack because it is possible to increase the(compatible) slack of the node y (which is 3 in FIG. 13C) by decreasingthe slack of node x (which is 7 in FIG. 13C) at the same time.

Examples of the Use of Sequential Optimization in a Design Flow

Referring again to FIG. 1, there is shown an illustrative diagram of thefirst design flow for the design of integrated circuit in accordancewith some embodiments of the invention. The following sections describesome details of constituent process modules of the design flow that usesequential optimization (SO) in accordance with principles of theinvention. It will be appreciated that the design flow of FIG. 1 is justone example design flow, and that the following sections describeexample design flow processes that employ SO in the course of theoverall design flow.

RTL Source Optimization Example

This section describes an example interactive RTL source optimizerprocess 104-1 of the design flow of FIG. 1. More specifically, thissection explains that one can utilize an interactive RTL editorincorporating sequential timing analysis to yield design improvements.

Consider the following example Verilog RTL source code:

module sample(in_a, in_b, in_c, in_d, in_e, in_f, in_g, in_h, clk);input [31:0] in_a; input [31:0] in_b; input [31:0] in_c; input [31:0]in_d; input [31:0] in_e; input [31:0] in_f; input [31:0] in_g; input[31:0] in_h; reg [31:0] reg_a; reg [31:0] reg_b; reg [31:0] reg_c;always @(posedge clk) begin  reg_a <= (((reg_b | in_a) & in_b) | in_c) &in_d;  reg_b <= reg_a | in_e;  reg_c <= ((reg_c | in_f) & in_g) | in_h;end endmodule

Assume synthesis does not modify the logical structure from what isgiven in the RTL, and suppose each logical operator requires the sameamount of time to evaluate (i.e. the gate delay for AND and OR gates arethe same).

A designer who is interested in improving the performance of this designmight focus on the large expression for reg_a, thinking that this is thecritical path of the design.

However, if the back-end synthesis tools are permitted to applysequential optimization techniques, this is not the true criticalsection of the design. reg_a and reg_b participate in a cycle, with atotal of 5 gate delays. Therefore clock skew scheduling and retiming canoptimize this cycle so that its minimum clock period is 2.5 gate delays.However, the cycle containing reg_c requires 3.0 gate delays toevaluate. Thus, if sequential optimization is taken into account, onewould better spend effort in optimizing this section of the designinstead.

Sequential optimization techniques, therefore, increases the flexibilityof the overall design process. In this example, the relative benefitthat can be achieved through sequential optimization techniques such asclock skew scheduling and retiming, is greater for the cycle containingregisters reg_a and reg_b than for the cycle containing register reg_c.Moreover, sequential optimization techniques can optimize the cycle withreg_a and reg_b to have a shorter minimum clock period (i.e. 2.5 agatedelays) than the cycle with reg_c (i.e. 3.0 gate delays). Sincesequential optimization techniques are able to shorten the minimum clockdelay for the cycle with reg_a and reg_b to be less than the minimumclock cycle with reg_c, it perhaps makes more sense to spend effortoptimizing the RTL source associated with the cycle containing reg_c.Under these circumstances, perhaps sequential optimization can be reliedupon to achieve improved minimum clock delay for the cycle containingreg_a and reg_b.

Integration of sequential timing analysis with the RTL editor can alsobe used to provide a mechanism for aiding manual path balancing. Forinstance, after running sequential timing analysis on the above design,the editor could make textual annotations to each of the three registerassignments indicating both their combinational delays and sequentialtiming information. The reg_a assignment would show a combinationaldelay of 4 and a sequential timing (e.g. the maximum mean cycle timeover all cycles containing the corresponding logic path) of 2.5. Thiswould inform the designer that this path best could be shortened throughsequential optimization, by moving logic across its register boundaries.The reg_c assignment would be given an annotation showing acombinational delay of 3 and sequential timing of 3, indicating to thedesigner that sequential optimization would not be effective, andsuggesting that RTL design optimization should be applied instead.

An RTL editor that uses sequential timing analysis to determinecriticality of design subsections would thus be able to highlightvisually (e.g. with a color change or some form of textual annotation)the appropriate portion of the RTL, to identify to the designer whereeffort should be spend for design optimization.

Data Path Optimization Example

This section describes an example data path optimization process 104-2of the design flow of FIG. 1. More specifically, this section explainsthat data path optimization can be affected by sequential slack data asthe following example illustrates.

FIGS. 14A-14B are illustrative drawings of a partial circuit designbefore (FIG. 14A) and after (FIG. 14B) data path optimization inaccordance with some embodiments of the invention. The partial designassumes gate delays are 1 and the MUX delay is 2. The cycle time of thepartial design is 4 (determined by the path from r₁ to r₂). However,register r₂ has some flexibility in its timing. If the data pathoptimization tool is made aware of the fact that signal a is on thecritical path and signal b is not, then it can make a combinationallogic transformation followed by a retiming to produce the circuit inFIG. 14B. The combinational transformation involves moving the signal aforward by making it the enable signal of the mux, and moving signal bbackward by removing it as the enable. The retiming moves register r2backwards across an AND gate. The cycle time in the circuit of FIG. 14Bis 3, and the design has been improved. This result cannot be achievedby iterating retiming and data path optimization, since in the case ofthis example partial circuit, neither technique applied alone would makeany changes to the original design.

Logic Optimization Example

This section describes an example logic optimization process 104-3 ofthe design flow of FIG. 1. More particularly, this section explains theuse of sequential slack for logic synthesis operations. The exampleshows that greater improvements can be achieved compared to a methodwhich uses only the combinational slack.

FIGS. 15A-15B are illustrative drawings of an example circuit before(FIG. 15A) and after (FIG. 15B) a possible logic optimization inaccordance with some embodiments of the invention. More specifically,FIG. 15A is an illustrative drawing of a circuit that is to be subjectto logic optimization process in accordance with principles of someaspects of the invention. It is assumed in this example that a path froma primary input A to a primary output B is critical and that the totaldelay of this path is larger than the target clock period. In thisexample we consider the following two possible operations to reduce thedelay of the path.

A first possible operation is to reduce the size of the inverter g1.This reduces the total capacitance driven by the NAND gate g2 and hencethe delay of the gate g2. However, this operation also increases thedelay of the inverter g1, and it is assumed that the total delay of gateg2 and gate g1 increases as well. However, decreasing the size of theinverter g1 reduces the overall area and power consumption of thecircuit.

A second possible operation is to directly connect the primary input Ato the NAND gate g3 and to connect the output of the combinational logicfrom the register g5 to gate g4 . The resulting circuit is shown in FIG.15B. The logic function of the circuit is not changed by this operation.The delay of the path from the primary input A to the primary output Bis reduced because the path is shorter and contains one NAND gate andone inverter less after the operation. The delay of the path from theregister g5 increases. The total area of the circuit is unchanged.

In this example, an automatic synthesis tool must to evaluate these twopossible operations: The tool has to check whether the operationimproves the worst slack and what the cost of the operation in area andpower is.

Both operations increase the slack of the critical path from the primaryinput A to the primary output B. The first operation increases the delayof the path leading towards register g5, and the second operationincreases the delay of the path leaving the register g5.

The combinational slack is computed on a register-to-register pathbasis. In this example the combinational slack is partitioned betweenthe two paths, the incoming and outgoing path of register g5: If theclock latency of register g5 is increased, the slack of the incomingpath increases, and the slack of the outgoing path decreases. If theclock latency of the register g5 is decreased, the effect is opposite.

If the combinational slack of the incoming path of register g5 is belowa certain limit, it might not be possible to perform the first operationwithout decreasing the worst slack. Similarly, if the slack of theoutgoing path is below a certain threshold, it might not be possible toperform the second operation without decreasing the worst slack. Theoptimization possibilities by the synthesis tool are limited and dependon the clock latency of register g5. Furthermore, if one operation isapplied it may no longer be possible to apply the second operation evenif the clock latency of register g5 is adjusted.

The view of the sequential slack analysis is different: each path has asequential slack equal to the sum of the combinational slacks of the twosub-paths; hence it does not depend on the clock latency of register g5.Thus, a sequential slack analysis looks at the slack of the entire cyclecontaining g5 and determines whether “sequential” slack is increased ordecreased by each of options 1 and 2—Whereas, traditional logicoptimization focuses separately on the combinational input sub-path tog5 and on the combinational output sub-path from g5. Both operations canbe evaluated and if both operations improve the worst slack, thesynthesis tool can chose the operation bringing the greatest improvementat the lowest cost.

Placement Example

This section describes an example sequential placement process 106-1 ofthe design flow of FIG. 1. The following example describes asequentially-driven placement methodology used to improve circuitperformance compared to traditional placement methods. Consider thefollowing netlist, given in standard Verilog format:

module test(A, B, C, D, X, Y); input A, B, C, D; output X, Y; NOR2G1(.A(A), .B(n4), .O(n1)); DFF G2(.D(n1), .Q(n2)); NOR2 G3(.A(n2),.B(C), .O(n3)); DFF G4(.D(n3), .Q(n4)); OR2 G5(.A(B), .B(n8), .O(n5));DFF G6(.D(n5), .Q(n6)); NAND2 G7(.A(n6), .B(D), .O(n7)); DFF G8(.D(n7),.Q(n8)); assign X = n4; assign Y = n8; endmoduleClock and power nets are omitted for simplicity.

Suppose, for example, that this design was constructed in a library inwhich all cells had a geometric height of 1 micron and a geometric widthof 1 micron, and for simplicity assume the pin locations for each celllie at the exact geometric center of the cell. Also suppose the designis to be placed in a bounding box with height of 4 microns and width 2microns, and suppose the inputs and outputs of the design have beenassigned fixed geometric locations. Specifically, assume that thelower-left corner of the placement bounding box is denoted as (0,0), andthe upper-right corner of the bounding box is denoted as (2,4).Moreover, the I/O locations are: A(0,0.5), B(0,1.5), C(2,3.5), D(2,2.5),X(0,3.5), Y(0,2.5)

FIG. 16A-16B are illustrative drawings of two candidate placements ofthe above circuit design. For clarity in the figures, nets are not drawnto connect to the geometric center of the cells where the pins shouldbe. However it is assumed that routing is done pin-to-pin to/from thecell centers in a Manhattan fashion.

Assume that the total wire length for the placement shown in FIG. 16A is15 microns, and the total wire length for the placement shown in FIG.16B is 17 microns. Thus a wirelength-driven placement tool will preferthe placement (that is, assign a lower cost to the placement) shown inFIG. 16A over the placement of FIG. 16B.

Suppose a simple delay model is given where each NAND2 gate has anintrinsic delay of 3 picoseconds, each NOR2 gate has an intrinsic delayof 5 picoseconds, each OR2 gate has an intrinsic delay of 7 picoseconds,and wires have a delay of 1 picosecond/micron. We allow setup and holdtimes of 0 picoseconds, and assume zero intrinsic delays for the DFFcells for simplicity. Under this timing model, the layout in FIG. 16Ahas a minimum clock period of 9 picoseconds (with the critical path fromG8(DFF) through G5(OR2) to G6(DFF)), while the layout of FIG. 16B has aminimum clock period of 10 picoseconds (also with a critical path fromG8 through G5 to G6, but with an extra picosecond of wire delay). Acombinational timing-driven placement tool would prefer the layout shownin FIG. 16A over that in FIG. 16B as well.

However, if we allow clock skew scheduling to take place afterplacement, the layout of FIG. 16B should have a higher figure of meritthan that of FIG. 16A. In FIG. 16A, the critical cycle is G1-G2-G3-G4,with a mean cycle time (total delay around the cycle divided by thenumber of registers) of 9 picoseconds. In FIG. 16B, the critical cycleis also G1-G2-G3-G4, but the mean cycle time here is 8 picoseconds (dueto decreased wire delay). Thus, a sequentially-driven placement tool canproperly identify the improvement in the maximum mean cycle time for thelayout in FIG. 16B and assign it a higher figure of merit than that ofFIG. 16A.

Post-Placement Retiming Example

This section describes an example sequential placement process 106-3 ofthe design flow of FIG. 1. The following example describes the use ofsequential analysis to improve the results of post-placement retimingoptimization.

Consider the following netlist, given in standard Verilog format:

module test(A, B, C, D, X); input A, B, C, D; output X; wire NB, ND, XD;INV I1(.A(B), .O(NB)); INV I2(.A(D), .O(ND)); OR4 O1(.A(A), .B(NB),.C(C), .D(ND), .O(XD)); DFF R1(.D(XD), .Q(X)); endmoduleClock and power nets are omitted for simplicity.

Assume that this design was constructed in a library where all cellshave a square geometry with height and width both 1 micron, and forsimplicity assume the pin locations for each cell lie at the exactgeometric center of the cell. Also suppose the design is to be placed ina square bounding box with height 3 microns and width 3 microns, andsuppose the inputs and outputs of the design have been assigned fixedgeometric locations. The lower-left corner of the placement bounding boxis denoted as (0,0). The upper-right corner of the bounding box isdenoted as (3,3). The I/O locations are: A(1.5,0), B(0,1.5), C(1.5,3),D(X(3,1.5).

Suppose this design has been placed as shown in FIG. 17A. For clarity inthe figures, nets are not drawn to connect to the geometric center ofthe cells where the pins should be. We assume routing is done pin-to-pinto/from the cell centers in a Manhattan fashion.

Suppose the delay of an INV gate is 1 picosecond and the delay of an OR4gate is 2 picoseconds. Assume registers have a 0 picosecond setuprequirement, and assume there is a wire delay of 1 picosecond per micronof wire.

Suppose the arrival time at all inputs is taken to be 0, and therequired arrival time at the output X is taken to be the same as theclock period of the design. Then the arrival time at all the inputs ofO1 is 2.5 picoseconds and the arrival time at the input of R1 is 5.5picoseconds. Therefore the clock period for this design without anychanges is 5.5 picoseconds.

Now suppose post-placement retiming optimization is allowed, but thatthis is done naively without regard to sequential analysis techniques.The register R1 would be retimed across the gate O1, in order to betterdistribute the slack on either side of R1. However, doing so wouldincrease the wire delay, regardless of how the retimed design is placed.To see why this must be the case, note that after the retiming therewill be a total of seven gates in the design (I1, I2, O1 and four newregisters introduced by the retiming of R1 across O1). This means thatat least one gate must lie in the uppermost third of the design (therectangle defined by the corners (0,2) and (3,3)). Therefore some wiremust cross from the bottom section (the rectangle defined by the corners(0,0) and (3,2)) to the top section. There must also be a wire back fromthe top section to the bottom section, as the output X lies in thebottom section. This additional meandering of wires will incur at leastan additional 2 units of delay somewhere, regardless of the layout. FIG.17B shows one possible layout for such a retimed design. Registersassociated with inputs A-D are named RA-RD, respectively. For thisplacement, the critical path is from RA to O1 to X, of length 4.5picoseconds. It can be shown that this placement is optimal for theretimed design with respect to the clock period.

While this naïve post-placement retiming improves the performance of thedesign, we can do much better by applying sequential timing analysis andoptimization. In this situation, the primary inputs and outputs can beconsidered a “host node” in the sequential timing graph. For theoriginal design in FIG. 17A, the critical cycle passes through the hostnode represented by input A, through O1, R1 and back to the host noderepresented by X. The critical cycle has a mean cycle time of 3picoseconds, meaning that sequential optimization can apply useful skewto the clock of R1 to obtain a clock period of 3 picoseconds. This is incontrast to the 4.5 picosecond clock period obtained by the retimingwithout consideration of sequential analysis.

The integration of the sequential analysis and optimization with thepost-placement retiming is essential. While one may certainly applyclock skew scheduling to the placement of FIG. 17B as a post-processingtechnique, this yields inferior results, as the performance improvementsdue to clock skew scheduling are insufficient to overcome thedegradation to the design made due to the extra wire delay added duringthe naïve post-placement retiming step.

Clock Tree Example

This section describes an example clock tree definition process 106-4 ofthe design flow of FIG. 1. The following example helps illustrate thebenefits that clock scheduling provides when the sequential slack isused to relax the accuracy required in creating the clock distributionnetwork. Assume two registers a and b are connected by two paths withdelay 10 and 20 units, respectively, forming a loop, as show in FIG. 18.The non-common path in the clock tree that leads to the two registersincludes three buffers, each inserting a delay of 2 units, for a totalof 6 units. Assume further that the target clock period is 22 timeunits. If, for simplicity, hold time violations are ignored, the designmust satisfy the following setup constraints:(2+2+2)(1+ε)+10≦22+(2+2+2)(1−ε)(2+2+2)(1+ε)+20≦22+(2+2+2)(1−ε)

By solving for ε we find that the largest permissible variation in theclock tree is

$ɛ = \frac{1}{6}$

The value ε denotes the permissible variation, or flexibility, of thedelays in the buffers of the clock tree, i.e., the delay of a buffervaries in the range 2(1−ε) to 2(1+ε). For simplicity, in this example weuse one value of ε for the entire tree, corresponding to the maximumpossible variation across all buffers. This technique is however notlimited to this case, and can be applied with different flexibilitiesfor each buffer in the tree. In addition, in this example we assume thatthe worst case variations in the delays of the paths connecting the tworegisters are already accounted for in their delays. An alternative thatcan be used in our approach is to consider the variations in theregister to register path delays explicitly.

The flexibility in this case is due to the 2 units of combinationalslack still available in the most critical of the two paths. From asequential point of view, however, the slack is much larger, and shouldtherefore lead to a larger flexibility in the clock tree. This can beseen by, for example, moving register b backwards in time by two units,i.e. by removing one of the buffers in the clock tree leading toregister b, thus altering the clock schedule. In this case, the setupconstraints become(2+2+2)(1+ε)+10≦22+(2+2)(1−ε)(2+2)(1+ε)+20≦22+(2+2+2)(1−ε)

By solving for ε again we find the largest permissible variation to be

$ɛ = \frac{2}{5}$which is much larger that the previous value of ⅙. An intermediateresult can be obtained by delaying the clock at register a by, forexample, adding an extra buffer with delay 2. In that case, theconstraints become(2+2+2+2)(1+ε)+10≦22+(2+2+2)(1−ε)(2+2+2)(1+ε)+20≦22+(2+2+2+2)(1−ε)giving a flexibility of

$ɛ = \frac{2}{7}$

An optimization technique can be employed to maximize the flexibilityavailable in the clock tree by using an appropriate clock schedule. Forexample, one optimization approach seeks to maximize designflexibilities. Given a circuit with a clock tree, additional delay canbe added at each branching point. A determination is made of how muchthe delay can increase or decrease for different points in the clocktree. The optimization problem is formulated as a linear program inwhich the variables are the flexibilities of the delays of each of thebuffers in the clock tree, the delay change at each branching point, andthe earliest and latest arrival times in the tree. The inequalityconstraints are the propagation of the earliest and latest arrival timesand the setup and hold constraints. The objective is to maximize theflexibilities. The optimization approach maximizes the minimumflexibility, then sets the flexibilities which form a cycle constraint,and then continues to maximize the next minimum flexibility.

Prototype of a Sequential Timing Analysis Tool

We have developed the sequential timing analysis tool (i.e. a computerprogram based process), that we have named “MAC”. Given a registertiming graph, it computes the critical cycle and the associatedlatencies required to achieve the minimum clock period. It can thenadjust the latencies to balance the slack, in accordance with thebalanced combinational slack process while maintaining the clock period.

For advanced features, the tool has the capability of distinguishingbetween different types of edges. This provides the opportunity to rundifferent optimization scenarios. In one particular optimization step ofthe scenario, the algorithm can be configured to ignore one subset ofthe edges, increase the slack of the edges of a second subset(parameterized edges), while ensuring that the slack of a third subsetdoes not decrease below a certain specified value (non-parameterizededges).

To illustrate the flexibility of this concept, consider an example withsetup and hold edges. Hold constraints are important if the clockschedule is implemented only by adjusting the latencies of the clocksignal at the registers and not by retiming. Hold time violations mayalso arise due to process variations in the clock tree. A hold timeviolation can be fixed by inserting additional delay with buffers or byrouting detours. Naturally, it is desirable to limit the number ofbuffers and routing detours.

Using the MAC computer program, the following optimization scenario ispossible. First, compute the minimum clock period ignoring all holdedges. Second, improve the slack of the setup edges further up to acertain specified value, after which the hold edges are added. Third,increase the slack of the hold edges subject to the constraint that theslack of a setup edge does not decrease if it is equal to or smallerthan the value up to which the slack was optimized. At this step, thehold edges are treated as parameterized edges (we want to increase theslack) whereas the setup edges are treated as non-parameterized edges(we just want to keep their slack). As explained above, these two setsof edges are handled differently by the balancing algorithm.

If for some reason it is not possible to insert buffers to fix hold timeviolations, another optimization scenario can be applied: first, onlythe hold edges are considered and the setup edges are ignored. Thebalance algorithm increases the slack of the hold edges up to zero, i.e.they are not violated. The purpose of this step is to check if all holdconstraints can be fulfilled. If not, only a limited number of all holdconstraints will be violated; the corresponding edges have negativeslacks. In a second optimization step, we can then add the setup edges,compute the minimum clock period subject to the hold constraints andimprove the slack on the setup edges.

Another aspect of the MAC program concerns the representation of thetiming data. The register timing graph as described with reference toFIGS. 6A-6B implies that the register-to-register delays are computedfor the entire circuit, which could be expensive. However is notnecessary to compute the entire graph. Using the capability todistinguish between parameterized and non-parameterized edge, thepresented algorithms can work directly on the gate-level timing graph.Thus, the principles of the invention are not limited to a registerlevel view of the circuit design, and they are applicable to a gatelevel—combinational logic level view—of a circuit design as well.

In this graph representation, each pin is represented by a vertex. Theedges correspond to the combinational timing arcs of the gates, thesetup and hold arcs of the registers, and the driver-to-load netconnections. In addition, in order to accurately model the differentdelays for rising and falling edges, one can introduce two nodes foreach pin, one node for the arrival time of the rising edge, and a secondnode for the arrival time of the falling edge. The size of theregister-to-register graph can potentially grow quadratically with thesize of the circuit, whereas the size of the gate-level timing graph isalways linear in the size of the circuit. Being able to work on thegate-level timing graph is an important feature allowing MAC to work onlarge designs.

Sequential Logic Synthesis

This section describes the application of sequential timing analysis inlogic synthesis. We first introduce the general concept behind ourprocedure and then we describe the experiments we performed using acomputer software tool known as, RTL Compiler, which is produced byCadence Design Systems of San Jose, Calif.

We motivate the use of sequential flexibility during logic synthesis byexamining statistics for the distribution of combinational slack acrossa typical design. FIG. 5.1 shows these statistics for one example. FIG.19A shows the distribution of slack over all the register-to-registerpaths in the design (paths with over 500 ps of slack were omitted). Ascan be seen, most paths are critical or near-critical in thecombinational sense. FIG. 19B shows the distribution of slacks for thesame circuit, after clock latencies have been determined for allregisters using the balanced combinational slack algorithm describedwith reference to FIGS. 10A-10D. Note the strong contrast between thetwo histograms: the vast majority of the paths which were initiallycritical have been given significant flexibility after slack balancing.

Noting the striking change to the combinational slack profile aftersequential flexibility is taken into account is the key to understandinghow exploiting sequential flexibility can bring about significantsavings in area, delay and power during logic synthesis. Withoutsequential flexibility, tools are forced to believe that most paths inthe design are critical, restricting the scope of optimizations that canbe performed. On the other hand, with the application of sequentialflexibility during optimization, synthesis tools have the freedom of anexpanded solution space, allowing substantial improvements to thedesign. This is best illustrated by noting how gate sizing algorithmstend to increase the size of gates on the critical path in order toensure timing requirements are met. In a typical synthesis flow, adesign such as the one in the example would be significantly over-sized,as many paths start off as being critical. However, by applying clocklatencies to distribute slack in the circuit, many paths are no longercritical, enabling a sequential synthesis tool to save large amounts ofarea.

In the next section, we describe experiments performed with RTL Compilerwhich quantify the benefits of using sequential flexibility during logicsynthesis.

Experiments with RTL Compiler

In the following we describe our experiments with RTL Compiler. We use arepeated adjustment of register clock latencies interleaved with thecombinational optimization capabilities of RTL Compiler to emulatesequential logic synthesis.

Performance Optimization

We implemented an optimization loop which iteratively extracts themaximum delay for each register-to-register path using the RTL Compilertiming analysis, computes new latencies and then optimizes the designwith RTL Compiler subject to the new synchronization latencies at theregisters. Algorithm 3 gives an outline of this optimization loop.

Algorithm 3 Sequential Logic Synthesis 1:  synthesize the circuit 2: repeat 3:   analyze timing and extract register timing graph 4:  sequential timing analysis and balance algorithm 5:   setsynchronization latencies 6:   combinational logic synthesis (eitherwith command ‘synthesize’      or command ‘synthesize - incremental’) 7: until terminated or no improvement

For the combinational optimization step we experimented with twodifferent options, implemented by the commands ‘synthesize’ and‘synthesize-incremental’. ‘synthesize-incremental’ performs mostly localoperations, for example adding or removing buffers and sizing orreplicating gates. It also globally optimizes the area of the design.The ‘synthesize’ command structures and restructures the design inaddition to the optimization mentioned before, and also performsredundancy removal.

FIG. 20 is an illustrative table that shows the characteristics of thedifferent sequential logic synthesis test cases and the results weobtained. The first five test cases are the largest test cases of theISCAS benchmark circuits and the last four test cases are customerdesigns. In this table ‘clock period’ is the minimum clock period withwhich the design can be run. The columns ‘LS’ show the clock period andarea which can be achieved in the standard RTL Compiler flow. Thefollowing columns show the result, which we obtained by running theoptimization loop (sequential optimization, logic synthesis) two times.The last column shows the relative change—a negative percentageindicates improvement. We remark that these are hypothetical results.The different latencies need to be realized by retiming or clockscheduling. Retiming may increase the number of registers and clockscheduling alone may not be possible if the difference in the latenciesare too large. Violations of hold constraints may have to be fixed byinserting additional delay buffers. For these test cases we decreasedthe clock period such that the slack was negative, and for most of thetest cases the slack remained negative even after the optimization. Theprimary goal was therefore to improve the cycle time. However, theresults also show that for some test cases the area could be decreasedquite substantially.

A detailed analysis has shown that the main improvements in the clockperiod are not due to the newly set synchronization latencies at theregisters, but rather that the clock period is decreased in thefollowing synthesis step. In some cases the worst path goes directlyfrom a primary input to a primary output without passing through anintermediate register. In this situation it is not possible to improvethe clock period by clock scheduling. However, one can considerassigning slack to the registers in the transitive fanin of the givenprimary output. This can enable refactorization to be done to reduce thedelay along the worst path and improve timing in this fashion.

Area Optimization

In the sequential optimization step we use the balance algorithm tocompute the latencies. It equally balances the slack such that the worstslack of all incoming paths equals the worst slack of all outgoingpaths. To reduce the area of the design this is obviously not the bestpossible solution. For area minimization, one needs to consider thearea-delay sensitivity of the incoming and outgoing logic cone of eachregister and adjust the register latency such that both sensitivitiesare equal R. Broderson, M. Horowitz, D. Markovic, B. Nikolic, and V.Stojanovic, “Methods for true power minimization,” in Digest ofTechnical Papers of the IEEE International Conference on Computer-AidedDesign, (San Jose, Calif.), pp. 35-42, November 2002. This is because ifthe sensitivity is unbalanced, one could simply gain area by increasingthe delay for the side with a higher area/delay ratio.

We implemented a simple heuristic algorithm that iteratively adjusts theregister latencies with the goal of reducing area. First, we increasethe latency of each register by the amount of the worst slack of allpaths leaving the register. Then we perform incremental synthesis.Similarly, we decrease the latencies of each register by the amount ofthe worst slack of all paths directed into the register, and againperform incremental synthesis.

Second, we implemented a loop over all registers: For each register weincrease or decrease the latency, do an incremental synthesis step forthe logic within the transitive fanin or transitive fanout, check theslacks and compute the change in area. Such a change in the latency isaccepted if the slack does not become negative and the area decreases.

With this additional optimization step we found we could decrease thearea further, by between 1.5% and 5.5%. Inspired by these results weintend to direct further research to developing a new global sequentialoptimization algorithm which can minimize the area subject to a givenclock period. We note that reducing area also reduces the powerconsumption of a chip, especially in the case that the latencyadjustments simply change the gate sizing.

Sequential Placement

Growing design sizes and increasingly aggressive timing requirementshave led to the abandonment of the traditional separation between logicsynthesis and physical optimization. Abstract modeling is no longersufficient to capture the true impact of physical effects on a finisheddesign. This holds equally true for our sequential synthesis flow.Retiming and clock skew scheduling must take interconnect delays intoaccount in a true and accurate fashion in order to be effective. Thus,in our studies we have included experiments that incorporate sequentialoptimization techniques with physical placement.

There are two primary concepts we wish to capture in combining layoutand sequential optimization. First, we wish to consider the impact ofwire delays on the overall design performance. Interconnect delay isgenerally an important consideration for those wires which lie on thecritical cycle. Introducing additional delay on the critical cycledirectly impacts the clock period of the final design, while thoseportions of the circuit which do not lie on the critical cycle cantypically tolerate larger delays. Second, we wish to account for theproblem of clock distribution to those registers which lie on thecritical cycle. We recognize that uncontrolled clock skew (e.g. throughprocess or manufacturing variation) is generally detrimental to theclock period of the final product when it is encountered on theregisters of the critical cycle, since there is zero timing flexibilityfor those registers. As such, we wish to minimize inadvertent skew forthe critical cycle.

Both of these primary concepts translate directly into spatial localityin the placement domain. Ensuring the critical cycle is localized to asmall area helps keep wirelengths short. Additionally, locality helpskeep the clock distribution network among the critical registers short,which helps reduce harmful skew. However, today's placement toolsgenerally do not deal with sequential optimization. Generally, themetrics in use today focus on minimization of total wirelength, withadditional consideration for combinational timing requirements. As such,there is a huge opportunity for optimization that is currently not beingexploited. To see this, we show in FIG. 6.1 placements obtained fromdifferent placement tools.

FIG. 21A shows a placement obtained using a traditional quadraticprogramming-based placement tool, similar to, GORDIAN. See, J. M.Kleinhans, G. Sigl, and F M. Johannes, “GORDIAN: A globaloptimization/rectangle dissection method for cell placement,” in Digestof Technical Papers of the IEEE International Conference onComputer-Aided Design, (Santa Clara, Calif.), pp. 506-509, November1988. FIG. 21B shows the same placement obtained using a combinationalslack-driven placer, using similar techniques to those found in moderntiming-driven placement tools. FIG. 21C shows the design placed usingCAPO, a leading-edge placement tool developed at UCLA. In all threeplacements, note that the critical cycle (shown darkened) spans across asubstantial portion of the die. The additional wire delay on thecritical cycle leads to significant degradation in performance. Also,the registers for the critical cycle are physically spaced apart, makingaccurate clock delivery much less reliable and hence subject todetrimental skew introduced by process and manufacturing variations. Forreference, the given design contains over 150,000 cells; the wirelengthsand distances involved are non-negligible.

In contrast, FIG. 21D shows a placement obtained using a prototypesequential placement tool we have developed, described below. Our toolhas clearly localized the critical cycle. This translates directly intoa faster design, due to less wireload on the critical portions, andgreater simplicity of clock tree construction necessary to support thefaster design. For this example, our tool yields a 50% reduction in wiredelay compared to the solution found by Capo, and a 77% reduction inwire delay compared to the combinational slack-driven placer.

We examined two approaches to introducing sequential optimizationawareness into a placement technique. The first experiment uses SoCEncounter, a computer software tool named produced by Cadence DesignSystems of San Jose Calif., which we augmented with our sequentialtiming analysis techniques. Results from this experiment are mixed, andsuggest that a placement tool that optimizes combinational timing doesnot adequately handle sequential flexibility, and thus such a tool maynot be ideal for incorporating into a sequential synthesis flow. For oursecond experiment, we developed a unique placement tool prototype. Thisnovel tool accounts for sequential flexibility using sequentialslack-based net weights along with physical layout constraints derivedfrom our sequential timing analysis. Although we must stress that thisis still an early prototype, the initial results obtained are excellent,suggesting that this will likely be a fruitful avenue for future work.

Experiments with SoC Encounter

Many concepts related to using sequential timing information duringplacement are explained in C. Hurst, P. Chong, A. Kuehlmann, PhysicalPlacement Driven by Sequential Timing Analysis, ICCAD, 2004, pp.379-386, November, 2004.

Similar to our experiments with RTL Compiler, we have implemented a TCLinterface between SoC Encounter and the clock scheduling program MAC toperform a sequential timing driven placement. The basic flow ispresented in Algorithm 4.

Algorithm 4 Sequential Placement with SOC Encounter 1:  Set latencies ofall registers to zero 2:  Timing driven placement (TDP 1) 3:  TrialRoute, extraction and timing analysis to extract the register    timinggraph 4:  Sequential timing analysis: Balanced Slack Algorithm (STA 1)5:  Set computed latencies at registers 6:  Timing driven placement (TDP2) 7:  Trial Route, extraction and timing analysis to extract theregister    timing graph 8:  Sequential timing analysis: Balanced SlackAlgorithm (STA 2)

First we generate a timing-driven placement with SoC Encounter. Based ona simple “Trial Route” (a procedure of Encounter) theregister-to-register timing information is extracted from the layout.New register latencies are computed using the balanced combinationalslack algorithm described above with reference to FIGS. 10A-10D. Astiming information comes from the initial placement, the latenciesshould thus be more accurate than a pre-placement timing estimationwithout wire delays. The computed clock latencies are applied to theregisters with the Encounter command set_clock_latency. With theselatencies a second timing-driven placement is generated with SoCEncounter. Finally, wire delays are extracted from this second placementand used to generate a second set of clock latencies.

The results are shown in Table 1. Our expectations were that the secondtiming driven placement generated with the clock latencies would givebetter results. These expectations were not fulfilled. In fact, for allour test cases we obtained exactly the opposite outcome: the slack ofthe new placement was worse.

TABLE 1 Clock TDP 1 CS 1 TDP 2 CS 2 Design Period (ns) Slack (ns) Slack(ns) Slack (ns} Slack (ns) Design A 7.50 −0.752 −0.467 −1.565 −0.442Design B 3.85 −3.053 −2.949 −4.378 −3.155 Design C 5.00 −32.173 −14.832−50.106 −37.787 Design D 1.43 −3.423 −2.653 −3.893 −2.720

Table 1 shows results obtained with SoC Encounter: First, a standardtiming driven placement is generated (TDP 1). Next, the timing isoptimized using the balanced combinational slack algorithm(CS 1). Basedon the new latencies another timing driven placement is generated (TDP2). Timing is optimized by applying a new clock schedule (CS 2).

FIGS. 22A-22D show a progression of placement results using theEncounter™ and the balanced combinational slack algorithm applied to oneexample circuit. FIG. 22A and FIG. 22C show critical paths. FIG. 22B andFIG. 22D show critical cycles. Cells are colored (shaded in B&W)according to module hierarchy. In FIG. 22B and FIG. 22D, the 20 mostcritical cycles are shown. As can be seen from the FIG. 22B and FIG.22D, there is little noticeable improvement in the physical locations ofcritical cycles found in the second placement, even after sequentialoptimization has been introduced. Although the experimental work usingthis approach is still fairly young, the results suggest that thisapproach using the Encounter placer is not entirely suited for dealingwith sequential criticality in a sequential synthesis flaw. We thus turnto a placement tool customized for our task.

Prototype of a Sequential Placement Tool

The following sections describe a prototype placement tool that usessequential timing information to maximize the potential ofpost-placement retiming or clock skew scheduling. The general procedure,outlined in Algorithm 5, involves three phases: (1) sequential timinganalysis, (2) assignment of weights based on sequential criticality, and(3) introduction of explicit cycle constraints.

Algorithm 5 Sequential Slack-Weighted Placement 1:  sequential timinganalysis 2:  assign net weights w(i) 3:  partition ^(P) ^(← allcells) 4: while ^(∃P(|P| > m)) do {GORDIAN main loop} 5:   solve globalconstrained QP 6:   bipartition all P where ^(|P| > m) 7:  solve finalglobal constrained QP 8:  (optional) do placement with cycle constraints(Algorithm 7) 9:  legalize placement into rows

We have implemented a modified version of GORDIAN See, J. M. Kleinhans,et al., “GORDIAN: A global optimization/rectangle dissection method forcell placement,”, supra. In this procedure, phases of globaloptimization are interleaved with bipartitioning. A quadraticprogramming (QP) problem is constructed to minimize the total weightedquadratic wirelengthΣα_(ij)[(x_(i)−x_(j))²+(y_(i)−y_(j))²]subject to a set of linear constraints. This problem is solved for theentire chip, and the positions of all cells are updated. Based on thisinformation, the cells in every subregion that contain more than mmembers are bipartitioned in a manner similar to that described by, R.S. Tsay, E. S. Kuh, and C.-P. Hsu, “PROUD: A sea-of-gates placementalgorithm,” IEEE Design & Test Of Computers, vol. 5, pp. 44-56, December1988, in order to minimize the total number of wires across the cut andmaintain reasonably balanced halves. The coordinates of the center ofeach subregion are computed and a linear center-of-gravity (COG)constraint is imposed on its members. The QP is updated to include thesenew constraints and the global optimization is repeated. GORDIAN is wellsuited to the requirements described above; nets can be easily weightedin both the global optimization and bipartitioning phases, andadditional constraints can be seamlessly imposed on the solution of theQP.

Sequential Slack Weighting

Net are assigned weights proportional to its relative sequentialcriticality. This is done to give priority to minimizing the lengths ofthe most critical wires as they are the most likely ones to limit theachievable clock period. After performing a sequential timing analysisas described above, we have a function S_(seq) that gives anapproximation of the sequential flexibility at each timing point; thisis the inverse of sequential criticality. A timing point is the input oroutput of any gate in the circuit; this is in contrast with the nodes inthe register timing graphs, which only correspond to the registers. Thesequential slack of a timing point is equal to the minimum slack for anyregister in its transitive fanin or fanout; i.e. the minimum slack ofany cycle in which that timing point participates.

We use the following equation to compute the net weight w(i):

${w(i)} = {1 + \frac{\beta}{\gamma + {{S_{seq}(i)}/\phi}}}$The constants β and γ are chosen to tune the distribution of weightsbetween the most and least critical nets. This is then applied to everyconnection α_(ij), in addition to scaling based on fanout.

This weighting alone is enough to produce layouts with improvedsequential timing characteristics, but its limitations should berecognized. Like their combinational counterparts, sequential slacks areinherently incompatible. Also, without computing the true sequentialslacks, the problems described above with respect to ordinary sequentialslack can also arise. Both of these problems can be solved with theintroduction of cycle constraints. Our iterative algorithm to handlethese constraints helps ensure that we catch all critical cycles.

Explicit Cycle Constraints

Assuming complete flexibility in assigning skew to all registers, for acycle

in the circuit to satisfy a target clock period φ, we must have,t _(g)

+t _(w)

≦N

φwhere t_(g)

is total intrinsic gate delay around l, t_(w)

is the total wireload delay around

, and N

is the number of registers in

.

Suppose we have a existing placement P′ in which the above constraint isviolated. Let t′_(w)

be the wireload delay around

for P′, and let d

be the total delay around

for P′. Then we have,

${{{t_{g}(\ell)} + {t_{w}^{\prime}(\ell)}} = {{{{d(\ell)}\mspace{14mu}{and}\mspace{14mu}\frac{t_{w}(\ell)}{t_{w}^{\prime}(\ell)}} \leq \frac{{{N(\ell)}\phi} - {t_{g}(\ell)}}{{d(\ell)} - {t_{g}(\ell)}}}\overset{\Delta}{=}{\mu(\ell)}}}\mspace{14mu}$which defines the wireload delay reduction factor μ

necessary for

to have a valid clock skew schedule for the target period.

Let (x′, y′) be the locations of the cells for the given placement P′.We wish to derive a new placement P=(x, y) which satisfies the abovegiven delay constraints. As an approximation, we take the wire delay fora cycle as being proportional to the sum of the squared Euclideandistances between cells in that cycle. That is,t _(w)

=ηΣ_((u,v)ε)

(x _(u) −x _(v))²+(y _(u) −y _(v))²where η is a constant. Thus the physical placement constraints are,

$\begin{matrix}{\frac{{\sum\limits_{{({u,v})} \in \ell}\;\left( {x_{u} - x_{v}} \right)^{2}} + \left( {y_{u} - y_{v}} \right)^{2}}{{\sum\limits_{{({u,v})} \in \ell}\;\left( {x_{u}^{\prime} - x_{v}^{\prime}} \right)^{2}} + \left( {y_{u}^{\prime} - y_{v}^{\prime}} \right)^{2}} \leq {\mu(\ell)}} & (6.1)\end{matrix}$

The denominator in inequality (6.1) as well as μ

are completely determined from the given placement and timinginformation. Thus inequality (6.1) contains only quadratic terms in (x,y). Also note that these constraints are convex.

We justify approximating total wire delay with the sum of squareEuclidean distances by our use of an iterative algorithm to solve theconstrained system. We aim to make only small changes to the layoutduring each iteration, so that any error in this approximation can besubsequently corrected. Details may be found below.

Lagrangian Relaxation

To realize the placement constraints, we use Lagrangian relaxation, astandard technique for converting constrained optimization problems intounconstrained problems. For brevity, we only present a simplifieddescription of this approach here. More information about Lagrangianrelaxation can be found in: P. M. Pardalos and M. G. Resende, eds.,Handbook of Applied Optimization. Oxford University Press, 2002; E.Golshtein and N. Treityakov, Modified Lagrangians and Monotone Maps inOptimization. John Wiley and Sons, 1996; and A. Srinivasan, K.Chaudhary, and E. S. Kuh, “RITUAL: A performance-driven placementalgorithm,” IEEE Transactions on Circuits and Systems, vol. 37, pp.825-839, November 1992.

Let f (x, y) be the sum of square wirelengths over all wires in thedesign for the placement (x, y). Recall that the classical analyticplacement formulation is simply the unconstrained problem min_(x,y) f(x,y). Our constrained problem then is

$\begin{matrix}{{\min\limits_{x,y}\;{{f\left( {x,y} \right)}\mspace{14mu}{such}\mspace{14mu}{that}\mspace{14mu}{g\left( {x,y} \right)}}} \leq 0} & (6.2)\end{matrix}$where the vector g represents the placement constraints. For each cyclein the design, there is a single element in g which corresponds to theconstraint inequality (6.1) for that cycle. We create the LagrangianL(x,y,k)=f (x,y)−k·g(x,y) where k is a vector of Lagrangian multipliers;k can be thought of as “penalties” which serve to increase the value ofthe cost function whenever a constraint is violated. The Lagrangian dualproblem is

$\begin{matrix}{\max\limits_{k \geq 0}{\min\limits_{x,y}{L\left( {x,y,k} \right)}}} & (6.3)\end{matrix}$

Our interest in the dual problem lies in the fact that, for convexproblems such as ours, a solution for (6.3) corresponds directly to asolution for the original problem (6.2). We use a standard technique ofsubgradient optimization to solve the dual (Algorithm 6).

Algorithm 6 Subgradient Optimization For Lagrangian Dual 1:  k ← 0 2: x,y ← arg min_(x,y) L(x,y,k) 3:  while KKT conditions are not satisfieddo 4:   k ← max(0,k+γ•g(x,y)) 5:   x,y ← arg min_(x,y) L(x,y,k)

For a fixed k, min_(x,y)L(x,y,k) is solved as an ordinary unconstrainedquadratic program. k is then adjusted based on the violated constraintswhich are found; if a constraint is violated, the corresponding penaltyis increased, so that subsequent iterations will reduce the violationsto minimize the cost function. Heuristics are available to determine anappropriate step size γ to adjust k; e.g. A. Srinivasan, K. Chaudhary,and E. S. Kuh, “RITUAL: A performance-driven placement algorithm,” IEEETransactions on Circuits and Systems, vol. 37, pp. 825-839, November1992; and P. M. Pardalos and M. G. Resende, eds., Handbook of AppliedOptimization. Oxford University Press, 2002. The Karush-Kuhn-Tucker(KKT) conditions for stopping the algorithm are described fully in P. M.Pardalos, et al., Handbook of Applied Optimization. Roughly speaking,the procedure stops once the penalty multipliers grow large enough toforce all constraint violations to zero.

Of course, the design may have many cycles, and thus there may be manyconstraints involved. We propose an iterative technique, given inAlgorithm 7, which reduces the number of cycles under consideration byignoring non-critical cycles.

Algorithm 7 Placement Using Cycle Constraints 1:  input: an initialplacement 2:  ^(T) ^(c) ^(←) current MMC, ^(S ←) {critical cycles} 3: while ^(T) ^(c) ^(> T) ^(f) do 4:   choose target clock period ^(T,T)^(f) ^(≦ T < T) ^(c) 5:   for all cycles ^(l ∈ S) do 6:    add cycleconstraint for ^(l) with target T 7:   remove all cells in S from COGbins 8:   solve QP with cycle constraints (ALGORITHM 6) 9:   reassignall cells in S to nearest COG 10:   solve QP with cycle constraints(ALGORITHM 6) 11:   ^(T) ^(c ←) current MMC, ^(S ← S ∪) {criticalcycles}

In each iteration, we add the critical cycles found in the currentplacement to the constraint set S. A clock period T is chosen which weuse as a target period for determining the cycle constraints for S. T isdecreased slowly from T_(c), the feasible clock period for the currentplacement, down to T_(f) , the final overall target clock period for thedesign. A slow adjustment of T helps ensure that we do not overconstrainthe current constraint set S while ignoring other cycles. That is, we donot wish to “squeeze too hard” on those cycles which are currentlycritical, as this may cause some other cycle not under consideration toviolate its timing constraint. Also, as noted before, we wish to perturbthe placement only by small amounts, so that any error in our quadraticapproximation of the wire delays can be corrected.

A significant benefit to using the iterative technique proposed inAlgorithm 7 is that we are able to correct errors in our estimate of thetrue sequential slack so that subsequent iterations may have a betterestimate of the sequential flexibility of each gate. Recall that truesequential slack S_(true) (v) is computationally expensive. Ourexperiments show balanced combinational slack and ordinary sequentialslack can be used to estimate true sequential slack in practicalapplications.

In this example approach, we use the ordinary sequential slack S_(seq)(v,v_(ref)) to estimate true sequential slack S_(true) (v); here wechoose v_(ref) to be a vertex on the critical cycle, in order tominimize the potential error. Nonetheless, there will still be errorpresent in this estimation. However, as we use an iterative techniquefor placement, the set of critical cycles tends to change with eachiteration, which helps to ensure that we compute S_(seq) with respect toseveral different choices of v_(ref). This way, if we mistakenlyidentify a critical vertex as non-critical, we will be able to correctthe mistake in subsequent iterations. In contrast, J. Cong and S. K.Lim, “Physical planning with retiming,” in Digest of Technical Papers ofthe IEEE/ACM International Conference on Computer-Aided Design, (SanJose, Calif.), pp. 1-7, November 2000, always takes v_(ref)=v_(ext), andso has no opportunity to correct such errors.

Center-of-gravity (COG) constraints are typically used in analyticplacement techniques to ensure that the cells are spread out relativelyevenly over the entire die area. We also wish our constrained placementto be appropriately spread out over the die area, but we do not wish theCOG constraints to overconstrain our solution. We approach this problemusing steps 7-10 in Algorithm 7, which allows critical cells to“migrate” to appropriate locations on the die to avoid violation oftiming constraints.

As a practical point, we also introduce cycles which are near-criticalduring each iteration, instead of only the critical cycles, to helpreduce the number of iterations performed. Instead of finding onecritical cycle, adding it to the constraints, solving the constrainedsystem to get new placement locations, and repeating this, adding onecycle at a time, we seek to add many cycles at a time, so that we reducethe number of times we update our constraints and compute newplacements, saving computation effort. But we should not add just anycycles, only those which we expect would become critical in lateriterations of this algorithm. As such, we only take the near-criticalcycles (i.e. cycles with very little slack) as these are the candidatesmost likely to become critical later on. Also, the main loop isterminated whenever either of the constrained QPs indicate that theproblem may have become overconstrained, as no further improvementbecomes possible in such case.

Our approach shares some similarity with that of, A. Srinivasan, K.Chaudhary, and E. S. Kuh, “RITUAL: A performance-driven placementalgorithm,” IEEE Transactions on Circuits and Systems, vol. 37, pp.825-839, November 1992, which also uses Lagrangian relaxation in ananalytic placement framework to resolve timing constraints. However,there are several key differences between our work and that of A.Srinivasan, K. Chaudhary et al. First, and most important, is that wedeal with the cyclic timing constraints which arise during clock skewscheduling, rather than simply path constraints. Second, we do not usethe analytic placement step itself to perform timing analysis. Thepractical effect of this is twofold: our approach allows us to usegeneral, nonlinear (and nonconvex) wire delay models, and we also do notencounter the degeneracy problems inherent in the constraints which comefrom timing analysis, as mentioned by A. Srinivasan, K. Chaudhary et al.Finally, we enjoy much greater computational efficiency, as ourLagrangian function can be seen as simply augmenting the weights ofedges between cells by k. Solving the Lagrangian dual for fixed krequires no more computation than solving an unconstrained QP for ourcircuit.

Row Legalization

Our approach to assigning cells to nonoverlapping positions in standardcell rows differs from the techniques typically used for this problem.Rather than the usual minimization of wirelength, our goal is instead tominimize the perturbation of the final layout with respect to thesolution of the QP, to avoid disturbing the placement of criticalcycles. To this end, we first sort all cells into the nearest cell rows,moving cells to adjacent rows whenever rows become overfilled. Next, foreach row we solve a linear program to minimize the sum of the deviationsof each cell from the QP solution. Linear constraints are added tohandle the non-overlap of cells within the row.

Experimental Setup

We ran our placement tool on a set of ten industrial benchmark circuitsas well as five of the largest synchronous designs from the ISCAS89benchmark suite. The ISCAS89 circuits were techmapped using anindustrial synthesis tool into an arbitrary library from the industrialbenchmarks.

The industrial libraries provided with the designs used interpolatedlookup-table based models to characterize the cells. Both capacitiveload and slew rate dependencies were incorporated in our timing model.The design technology files gave the electrical characterization for thewires; in all cases, we assumed the use of metal layer 3 for routing. Weused the half-perimeter bounding box metric as our estimate of thewirelength, noting that our algorithms are actually independent of thewireload estimation technique used, unlike other works, e.g. A.Srinivasan, K. Chaudhary, et al., supra.

Currently, our placement tool can only handle single-row cells, so wefound it necessary to convert some circuit elements to single-rowinstances, for the purpose of our experiments. Double-row cells weregiven a different aspect ratio, keeping the same area. Large macros weregiven an arbitrary size. I/O pads were assigned randomly around the dieperimeter.

Limitations in our timing analysis tool required some design changes tobe made. Transparent latches were treated as ordinary registers, andcombinational cycles were broken arbitrarily. Some hard macros did nothave timing information associated with them, so for the purpose oftiming analysis hard macros were treated as if they were I/Os for theoverall circuit. Some designs used multiple clock domains. As we had noadditional information regarding the relative phases and clockfrequencies of such, we uniformly regarded the circuits as having only asingle clock domain. We note, however, that the techniques described inthis paper can be easily extended to multiple clock domains.

Experimental Results

Our experimental results are shown in table of FIG. 23. The Size columnindicates the number of placed instances for the design. The Reg columnindicates the number of registers. The NW MMC column gives the MMC forthe design when no wireload is taken into consideration; this is theminimum feasible clock period. The REG MMC column shows the MMC achievedfor a completely placed design using our placement flow with equalweights attached to the wires; effectively, this is a placement toolsimilar to GORDIAN. The COM MMC column shows the MMC achieved afterplacement using a combinational slack-based weighting function for thenets.

The SEQ MMC column indicates the MMC achieved after placement using thesequential slack-based weighting for the nets. The percentage figureindicates the wire delay reduction for the SEQ MMC result compared withthe COM MMC result. This is arguably a better figure of merit than theabsolute reduction in clock period, since no placer can ever hope toreduce the clock period below the no wireload MMC. The Run columnindicates the run time for this algorithm, in seconds.

The CYCLE MMC column indicates the MMC achieved after placement usingthe cycle constraint technique described above, again with thepercentage indicating reduction in wire delay compared to thecombinational-weighted technique, and the Run column indicating the runtime in seconds.

We also compare our placement tool against Capo, a leading-edge placerwhich focuses on wirelength minimization. See A. E. Caldwell, A. B.Kahng, and I. L. Markov, “Can recursive bisection alone produce routableplacements?,” in ACM/IEEE Design Automation Conference, pp. 477-482,2000. As the two placers have different objectives, we certainly do notexpect either one to be competitive in the other's problem domain.However, the comparison serves to quantity the benefit of using asequential flexibility-aware placer, rather than choosing an placerwhich is best suited for another task. The CAPO MMC column in the tableof FIG. 23 shows the MMC obtained after placement using Capo, the Runcolumn indicates the run time for Capo in seconds, and the CYvsCA columnindicates the percentage improvement in wire delay of our cycleconstraint-based technique compared to the placement from Capo.

We show significant improvement in achievable clock period throughapplication of our algorithm. We achieved an improvement in wire delayof 27.7% over a combinational slack-weighted placement technique, and32.2% improvement over the results of Capo.

On average, our tool gives a 57.1% increase in total wirelength comparedwith Capo. We note that our tool currently does little to control thetotal wirelength of the final placed design. As there are many nets inthe design which are not critical, there is much opportunity for us tofurther reduce wirelength, especially during row legalization, which, asnoted before, seeks only to minimize the perturbation of the QP solutionand ignores wirelength completely. Additionally, modern partitioningtechniques such as those found in Capo can replace the simplepartitioning algorithm we use, especially at the coarsest levels ofpartitioning, where the information provided by the QP solution islimited. These are just a couple of the techniques we can apply tofurther reduce wirelength in our tool.

Run times are measured in CPU seconds on a 3 GHz Pentium 4 processor.Our placement flow is competitive with Capo with regard to run time,even with the inclusion of computationally expensive sequential timinganalysis which is absent in Capo. Still, there is significant potentialto decrease run time further.

Post-Placement Retiming and Clock Skew Scheduling General Concept

Clock skew scheduling and retiming are two alternative techniques toimplement the synchronization latencies assumed for the sequentialoptimization steps described before. In clock skew scheduling, theintentional differences in the clock arrival times, also referred to as“useful skew”, are implemented by designing dedicated delays into theclock distribution. Alternatively, retiming balances the path delays byrelocating the registers.

Retiming, applied as post-placement step, is a straightforward extensionof existing in-place optimization techniques such as re-buffering orlocal logic restructuring. See, O. Coudert, J. Cong, S. Malik, and M. S.Sarrafzadeh, “Incremental cad,” in Digest of Technical Papers of theIEEE/ACM International Conference on Computer-Aided Design, (San Jose,Calif.), pp. 236-243, November 2000. For a given placement, e.g. derivedby the algorithm described above, the exact timing of the gates andfairly good estimates of the interconnect timing are known. One possibleapproach to compute the new register positions is based on thetraditional delay-constrained min-area retiming formulation. See, I. S.Kourtev and E. G. Friedman, Timing Optimization Through Clock SkewScheduling. Norwell, Mass.: Kluwer Academic Publishers, 2001. Forimplementing the retiming solution, the placement should be updatedincrementally to provide space for the new registers. This, in turn,changes some of the interconnect delays thus possibly requiring arepeated application of retiming and incremental placement until itreaches a stable solution. Ultimately, the register retiming should beintegrated into the second phase of the sequential placement algorithmdescribed above. By considering the register areas during the iterativeprocessing of the critical cycle, the delay updates are automaticallytaken into account and a legal placement is generated.

In a combined approach of retiming and clock skew scheduling, theavailability of dedicated clock latencies can be used to relax thestrict requirements on register relocations. For example, instead ofmoving the register backward through a gate, its clock latency can beincreased resulting in an identical impact on the timing behavior.Depending on the range of implementable clock latencies, the applicationof retiming can then be reduced to the few cases where the clock delayswould be too large for a reliable implementation.

Clearly, there is a three way trade-off between (1) the amount ofregister retiming moves, (2) the capability of reliably implementinglarge clock latencies for individual registers, and (3) the amount ofsequential flexibility usable for sequential optimization. For example,in the absence of retiming and with a zero-skew clock distribution, theoptimization potential is reduced to the pure combinational case. Ifonly retiming is available, a large number of retiming moves might berequired to implement substantial sequential optimization moves. This,in turn, may increase the number of registers, power consumption, andresources for clock routing. On the other hand, if large clock latenciescan be implemented in a reliable manner, retiming might be fullyavoided, addressing some of the practical acceptance problems. Inanother extreme, the use of large clock latencies could enable alow-power design methodology based on min-area retiming.

Overview of Current Clock Implementation Methods

Traditionally, clock distribution networks were implemented using a“zero-skew” methodology, which minimizes the difference of the clocksignal arrival time at the registers. Such approach directly reflectedthe underlying combinational synthesis paradigm where the longest delayof any combinational path adjusted by the unintentional clock skewbetween its two registers determines the maximal clock frequency of adesign. High-performance designs with extreme requirements on low clockskew use expensive clock distribution methods such as clock meshes,H-trees, or a combination. See, P. Restle, T. McNamara, D. Webber, P.Camporese, K. Eng, K. Jenkins, D. Allen, M. Rohn, M. Quaranta, D.Boerstler, C. Alpert, C. Carter, R. Bailey, J. Petrovick, B. Krauter,and B. McCredie, “A clock distribution network for microprocessors,”Journal of Solid-State Circuits, vol. 36, no. 5, pp. 792-799, 2001.Because of their high power consumption and large routing resources suchmethods are only practical for selected custom-designed chips forcritical parts of ASIC designs.

Low-cost clock solutions for ASIC designs are typically based on abottom-up construction of a balanced or unbalanced clock tree using somemodification of the clock tree synthesis algorithm presented in R. S.Tsay, “An exact zero-skew clock routing algorithm,” IEEE Transactions onComputer-Aided Design of Integrated Circuits and Systems, vol. 12, pp.242-249, February 1993. In the past few years, clock skew scheduling hasfound some adaptation in ASIC design flows. See, C. Leiserson and J.Saxe, “Optimizing synchronous systems,” Journal of VLSI and ComputerSystems, vol. 1, pp. 41-67, January 1983; I. S. Kourtev and E. G.Friedman, Timing Optimization Through Clock Skew Scheduling. Norwell,Mass.: Kluwer Academic Publishers, 2001; C. W. Tsao and C.-K. Koh,“UST/DME: a clock tree router for general skew constraints,” in Digestof Technical Papers of the IEEE/ACM International Conference onComputer-Aided Design, (San Jose, Calif.), pp. 400-405, November 2000;and J. Xi and W.-M. Dai, “Useful-skew clock routing with gate sizing forlow power design,” in Proceedings of the 33rd ACM/IEEE Design AutomationConference, (Las Vegas, Nev.), pp. 383-388, June 1996. In theseapplications, clock skew scheduling is applied as a mere post-placementoptimization step for improving the cycle time. Most practical toolswith the exception of, S. Held, B. Korte, J. Maβberg, M. Ringe, and J.Vygen, “Clock scheduling and clocktree construction for high-performanceASICs,” in, Digest of Technical Papers of the IEEE/ACM InternationalConference on Computer-Aided Design, (San Jose, Calif.), pp. 232-239,November 2003, first compute a single clock schedule which is thenimplemented by the clock tree synthesis step. In S. Held et al., a clocktree synthesizer is presented that accepts ranges of clock latencyvalues provided by the scheduler.

Up until now, clock tree synthesis and clock scheduling have beenimplemented as two disconnected steps in the design flow. The clockscheduler computes a set of register clock latencies that optimizes theperformance for a given design. The computed latencies or intervals ofthe latencies are implemented by the clock tree synthesizer. Ultimately,both the scheduling step and clock tree synthesis should be combined forglobally optimizing the clock tree topology and clock schedule.

Multi-Domain Clock Skew Scheduling

In practice, a clock schedule with large variations of the registerlatencies generally cannot be realized in a reliable manner. This isbecause the implementation of dedicated delays using additional buffersand interconnections is highly susceptible to within-die variations ofprocess parameters. As a consequence, the practically applicable maximumdifferences for the clock arrival times are typically restricted to lessthan 10% of the clock period, which limits the optimization potential ofclock skew scheduling.

Multiple clocking domains are routinely applied in designs to realizeseveral clocking frequencies and also to address specific timingrequirements. For example, a special clocking domain that delivers aphase-shifted clock signal to the registers close to the chip inputs andoutputs is regularly used to achieve timing closure for ports withextreme constraints on their arrival and required times.

A multi-domain approach could also be used to realize larger clocklatency variations for all registers. In combination with awithin-domain clock skew scheduling algorithm, they could implement anaggressive sequential optimization that would be impractical withindividual delays of register clocks. The motivation behind thisapproach is based on the fact that large phase shifts between clockingdomains can be implemented reliably by using dedicated, possiblyexpensive circuit components such as “structured clock buffers, ” K. M.Carrig, “Chip clocking effect on performance for IBM's SA-27E ASICtechnology,” IBM Micronews, vol. 6, no. 3, pp. 12-16, 2000, adjustmentsto the PLL circuitry, or simply by deriving the set of phase-shifteddomains from a higher frequency clock using different tapping points ofa shift register.

In current design methodologies, the specification of multiple clockingdomains is mostly done manually as no design automation support isavailable. We have developed a new algorithm for constrained clock skewscheduling which computes for a user-given number of clocking domainsthe optimal phase shifts for the domain clocks and the assignment of thecircuit registers to the domains. See, K. Ravindran, A. Kuehlmann, andE. Sentovich, “Multi-domain clock skew scheduling,” in Digest ofTechnical Papers of the IEEE/ACM International Conference onComputer-Aided Design, (San Jose, Calif.), pp. 801-808, November 2003.

For the clock distribution within a domain, the algorithm can assume azero-skew clock delivery or apply a user-provided upper bound for thewithin-domain latency. Experiments demonstrate that a clock skewschedule using a few domains combined with a small within-domain latencycan reliably implement the full optimization potential of anunconstrained clock schedule.

Our algorithm is based on a branch-and-bound search for the assignmentof registers to clocking domains. For every complete assignment ofregisters to domains, the timing graph is analyzed and the best cycletime is improved if possible. We apply a satisfiability (SAT) solverbased on a problem encoding in conjunctive normal form (CNF) toefficiently drive the search and compactly record parts of the solutionspace that are guaranteed to contain no solutions better than thecurrent one. The combination of a modern SAT solver, M. W. Moskewicz, C.F. Madigan, Y. Zhao, L. Zhang, and S. Malik, “Chaff: Engineering anefficient SAT solver,” in Proceedings of the 38th ACM/IEEE DesignAutomation Conference, (Las Vegas, Nev.), pp. 530-535, June 2001, withan underlying orthogonal optimization problem provides a powerfulmechanism for a hybrid search that has significant potential for otherapplications in many domains. More details of the algorithm can be foundin, K. Ravindran et al., in Digest of Technical Papers of the IEEE/ACMInternational Conference on Computer-Aided Design, supra.

Experiments indicate that despite the potential complexity of theenumeration process, the multi-domain clock skew scheduling algorithm isefficient for modestly sized circuits and works reasonably fast forcircuits with several thousand registers. Our results show that aconstrained clock skew schedule with few clocking domains and zero or 5%within-domain latency can in most cases achieve the optimal cycle timedictated by the critical cycle of the circuit. We believe thatmulti-domain clock skew scheduling provides an important component of anoverall clocking solution for the critical design parts.

FIG. 24 is an illustrative block level diagram of a computer system 2400that can be programmed to implement processes involved with theoptimization of circuit design using sequential timing information inaccordance with embodiments of the invention. Computer system 2400 caninclude one or more processors, such as a processor 2402. Processor 2402can be implemented using a general or special purpose processing enginesuch as, for example, a microprocessor, controller or other controllogic. In the example illustrated in FIG. 24, processor 2402 isconnected to a bus 2404 or other communication medium.

Computing system 2400 also can include a main memory 2406, preferablyrandom access memory (RAM) or other dynamic memory, for storinginformation and instructions to be executed by processor 2402. Mainmemory 2406 also may be used for storing temporary variables or otherintermediate information during execution of instructions to be executedby processor 2402. Computer system 2400 can likewise include a read onlymemory (“ROM”) or other static storage device coupled to bus 2404 forstoring static information and instructions for processor 2402.

The computer system 2400 can also include information storage mechanism2408, which can include, for example, a media drive 2410 and a removablestorage interface 2412. The media drive 2410 can include a drive orother mechanism to support fixed or removable storage media 2414. Forexample, a hard disk drive, a floppy disk drive, a magnetic tape drive,an optical disk drive, a CD or DVD drive (R or RW), or other removableor fixed media drive. Storage media 2414, can include, for example, ahard disk, a floppy disk, magnetic tape, optical disk, a CD or DVD, orother fixed or removable medium that is read by and written to by mediadrive 2410. Information storage mechanism 208 also may include aremovable storage unit 2416 in communication with interface 2412.Examples of such removable storage unit 2416 can include a programcartridge and cartridge interface, a removable memory (for example, aflash memory or other removable memory module). As these examplesillustrate, the storage media 2414 can include a computer useablestorage medium having stored therein particular computer software ordata.

In this document, the terms “computer program medium” and “computeruseable medium” are used to generally refer to media such as, forexample, memory 2406, storage device 2408, a hard disk installed in harddisk drive 210. These and other various forms of computer useable mediamay be involved in carrying one or more sequences of one or moreinstructions to processor 2402 for execution. Such instructions,generally referred to as “computer program code” (which may be groupedin the form of computer programs or other groupings), when executed,enable the computing system 2400 to perform features or functions of thepresent invention as discussed herein.

The foregoing description and drawings of preferred embodiment inaccordance with the present invention are merely illustrative of theprinciples of this invention, and that various modifications can be madeby those skilled in the art without departing from the scope and spiritof the invention.

The invention claimed is:
 1. A circuit design placement retiming methodcomprising: using a computer system to run a process to perform thesteps of: determining at least first and second placement registerplacement alternatives; determining for the first register placementalternative, a first value indicative of a proportionality of delay tonumber of registers for a structural cycle of the first registerplacement alternative having a largest proportionality of delay tonumber of registers in the first register alternative; determining forthe second register placement alternative, a second value indicative ofa proportionality of delay to number of registers for a structural cycleof the second register placement alternative having a largestproportionality of delay to number of registers in the second registeralternative; using the determined first and second values to evaluatethe first and second register placement alternatives; selecting one ofthe first and second register placement alternatives based upon theevaluation; and incorporating the selected placement alternative intothe circuit design.
 2. The method of claim 1, wherein selecting involvesselecting the alternative having a value indicative of the lowerproportionality of delay to number of registers.
 3. An article ofmanufacture that includes a computer readable storage medium encodedwith program code to cause a computer system to perform a process thatcomprises: using the computer system to run a process to perform thesteps of: determining at least first and second placement registerplacement alternatives; determining for the first register placementalternative, a first value indicative of a proportionality of delay tonumber of registers for a structural cycle of the first registerplacement alternative having a largest proportionality of delay tonumber of registers in the first register alternative; determining forthe second register placement alternative, a second value indicative ofa proportionality of delay to number of registers for a structural cycleof the second register placement alternative having a largestproportionality of delay to number of registers in the second registeralternative; using the determined first and second values to evaluatethe first and second register placement alternatives; selecting one ofthe first and second register placement alternatives based upon theevaluation and incorporating the selected placement alternative into thecircuit design.
 4. The article of manufacture of claim 3, whereinselecting involves selecting the alternative having a value indicativeof the lower proportionality of delay to number of registers.